Hybrid tracking of transaction read and write sets

ABSTRACT

Embodiments of the invention relate to tracking processor transactional read and write sets, thereby eliminating speculative mis-predictions. Both non-speculative read set and write set indications are maintained for a transaction. The indications are stored in cache. In addition, load and write queues of addresses are maintained. The load queue of addresses relates to speculative members of a read set and the write queue of addresses relates to speculating member of a write set. For a received read request, a transaction resolution process takes place, and a resolution is performed if an address match in the write queue is detected. Similarly, for a receive write request the transaction interference additionally checks the load queue and the non-speculative read set for the pending address.

BACKGROUND

The present embodiments relate to transactional execution and trackingof memory data. More specifically, the embodiments relate to trackingprocessor transactional read and write sets to eliminate speculativemis-predictions.

The number of central processing unit (CPU) cores on a chip and thenumber of CPU cores connected to a shared memory continues to growsignificantly to support growing workload capacity demand. Theincreasing number of CPUs cooperating to process the same workloads putsa significant burden on software scalability. For example, shared queuesor data-structures protected by traditional semaphores become hot spotsand lead to sub-linear n-way scaling curves. Traditionally this has beencountered by implementing finer-grained locking in software, and withlower latency/higher bandwidth interconnects in hardware. Implementingfine-grained locking to improve software scalability can be verycomplicated and error-prone, and at today's CPU frequencies, thelatencies of hardware interconnects is limited by the physical dimensionof the chips and systems, and by the speed of light.

Implementations of hardware Transactional Memory (TM) have beenintroduced, wherein a group of instructions, called a transaction,operate atomically and in isolation (sometimes called “serializability”)on a data structure in memory. The transaction is a sequence ofinstructions that appears as if they have all been executed without anyintervening interaction with another processor. The transaction executesoptimistically without obtaining a lock, but may need to abort and retrythe transaction execution if an operation, of the executing transaction,on a memory location conflicts with anther operation on the same memorylocation, also referred to as interference. Instructions are groupedtogether, and transactional memory requires the tracking or memory databeing used. Tracking generally occurs as a cache line granularity. Twoseparate sets are tracking, including a read set and a write set. Theread set includes all cache lines that have been read by a currenttransaction. The write set includes all cache lines that have been readby the current transaction.

SUMMARY

A computer program product and system are provided for trackingprocessor transactional read and write sets.

In one aspect, a computer program product is provided to addressover-indication for both a non-speculative read set and anon-speculative write set. Two non-speculative sets of data aremaintained, including a non-speculative read set indication ofnon-speculative reads of data stored in a cache by a processor cacheunit for a transaction by a first requestor and a non-speculative writeset indication of non-speculative writes of written data stored in thecache by the processor cache unit by the first requestor. A first queueof addresses is maintained corresponding to speculatively executedmemory read instructions corresponding to speculative members of a readset, and a second queue of addresses is maintained corresponding tospeculatively executed memory write instructions corresponding tospeculative members of a write set. A transaction interferenceresolution is performed in response receiving a request for data by aremote processor. The resolution determines a potential transactioninterference exclusively conflicting with a speculative instruction andholds a response to the received request until the speculativeinstruction is committed or flushed.

In a further aspect, a computer system is provided with a processingunit operatively coupled to memory. A tool is provided in communicationwith the processing unit. The tool functions to track processortransactional read and write sets. Both a non-speculative read set and anon-speculative write set are maintained. The non-speculative read setis an indication of non-speculative reads of data stored in a cache by aprocessor cache unit for a transaction by a first requestor. Thenon-speculative write set is an indication of non-speculative writes ofwritten data stored in the cache by the processor cache unit by thefirst requestor. Two queues of addresses are maintained, including afirst queue and a second queue. The first queue of addresses correspondsto speculatively executed load instructions corresponding to speculativemembers of a read set, and the second queue of addresses correspondingto speculatively executed store instructions corresponding tospeculative members of a write set. The tool performs a transactioninterference resolution in response to receipt of a request for data bya remote processor. The transaction interference includes ascertainingif there is a potential transaction interference that exclusivelyconflicts with a speculative instruction. The tool holds a response tothe received request until the speculative instruction is eithercommitted or flushed.

These and other features and advantages will become apparent from thefollowing detailed description of the presently preferred embodiment(s),taken in conjunction with the accompanying drawings.

BRIEF DESCRIPTION OF THE SEVERAL VIEWS OF THE DRAWINGS

The drawings reference herein forms a part of the specification.Features shown in the drawings are meant as illustrative of only someembodiments, and not of all embodiments unless otherwise explicitlyindicated.

FIG. 1 depicts a flow chart illustrating a process for determining apotential conflict associated with execution of transactions on at leasttwo processors in a transactional memory system.

FIG. 2 depicts a flow chart illustrating the flow for instructionsthrough a microprocessor while managing the speculative access queue(s).

FIG. 3 depicts a block diagram illustrating a pending access queue.

FIG. 4 depicts a set of flow charts illustrating methods performed inconjunction with management of a pending address queue.

FIG. 5 depicts a flow chart illustrating steps for managing the pendingqueue, and specifically performing interference checks with respect toaddresses corresponding to received requests from remote processors.

FIG. 6 depicts a flow chart illustrating management of the queue, andmore specifically, addressing overflow of the queue.

FIG. 7 depicts a flow chart illustrating another form of management ofthe pending address queue that does not stall the update of the queue.

FIGS. 8A-8C depict a flow chart illustrating a process for holding aresponse to a requesting processor based on potential interferenceassociated with a transaction.

FIGS. 9A-9C depicts a flow chart illustrating another process forholding a response to a requesting processor based on potentialinterference associated with a transaction.

FIG. 10 depicts a block diagram showing components of an example CPU, inaccordance with the present embodiments.

FIG. 11 depicts a block diagram showing an example multicoretransactional memory environment, in accordance with an illustrativeembodiment.

FIG. 12 depicts a block diagram showing an example multicoretransactional memory environment, in accordance with an illustrativeembodiment;

DETAILED DESCRIPTION

It will be readily understood that the components of the presentembodiment(s), as generally described and illustrated in the Figuresherein, may be arranged and designed in a wide variety of differentconfigurations. Thus, the following detailed description of theembodiments of the apparatus, system, and method of the presentembodiment(s), as presented in the Figures, is not intended to limit thescope of the embodiment(s), as claimed, but is merely representative ofselected embodiments.

Reference throughout this specification to “a select embodiment,” “oneembodiment,” or “an embodiment” means that a particular feature,structure, or characteristic described in connection with the embodimentis included in at least one embodiment described herein. Thus,appearances of the phrases “a select embodiment,” “in one embodiment,”or “in an embodiment” in various places throughout this specificationare not necessarily referring to the same embodiment.

The illustrated embodiments will be best understood by reference to thedrawings, wherein like parts are designated by like numerals throughout.The following description is intended only by way of example, and simplyillustrates certain selected embodiments of devices, systems, andprocesses that are consistent with the embodiment(s) as claimed herein.

Historically, a computer system or processor had only a single processor(also known as a processing unit or central processing unit). Theprocessor included an instruction processing unit (IPU), a branch unit,a memory control unit and the like. Such processors were capable ofexecuting a single thread of a program at a time. Operating systems weredeveloped that could time-share a processor by dispatching a program tobe executed on the processor for a period of time, and then dispatchinganother program to be executed on the processor for another period oftime.

As technology evolved, memory subsystem caches were often added to theprocessor as well as complex dynamic address translation includingtranslation lookaside buffers (TLBs). The IPU itself was often referredto as a processor. As technology continues to evolve, an entireprocessor could be packaged as a single semiconductor chip or die, sucha processor was referred to as a microprocessor. Then processors weredeveloped that incorporated multiple IPUs, such processors were oftenreferred to as multi-processors. Each such processor of amulti-processor computer system (processor) may include individual orshared caches, memory interfaces, system bus, address translationmechanism and the like. Virtual machine and instruction set architecture(ISA) emulators added a layer of software to a processor, that providedthe virtual machine with multiple “virtual processors” (aka processors)by time-slice usage of a single IPU in a single hardware processor. Astechnology further evolved, multi-threaded processors were developedenabling a single hardware processor having a single multi-thread IPU toprovide a capability of simultaneously executing threads of differentprograms, thus each thread of a multi-threaded processor appeared to theoperating system as a processor. As technology further evolved, it waspossible to put multiple processors (each having an IPU) on a singlesemiconductor chip or die. These processors were referred to processorcores or just cores. Thus the terms such as processor, centralprocessing unit, processing unit, microprocessor, core, processor core,processor thread, and thread, for example, are often usedinterchangeably. Aspects of embodiments herein may be practiced by anyor all processors including those shown supra, without departing fromthe teachings herein. Wherein the term “thread” or “processor thread” isused herein, it is expected that particular advantage of the embodimentmay be had in a processor thread implementation.

In “Intel® Architecture Instruction Set Extensions ProgrammingReference” 319433-012A, February 2012, incorporated herein by referencein its entirety, Chapter 8 teaches, in part, that multithreadedapplications may take advantage of increasing numbers of CPU cores toachieve higher performance. However, the writing of multi-threadedapplications requires programmers to understand and take into accountdata sharing among the multiple threads. Access to shared data typicallyrequires synchronization mechanisms. These synchronization mechanismsare used to ensure that multiple threads update shared data byserializing operations that are applied to the shared data, oftenthrough the use of a critical section that is protected by a lock. Sinceserialization limits concurrency, programmers try to limit the overheaddue to synchronization.

Intel® Transactional Synchronization Extensions (Intel® TSX) allow aprocessor to dynamically determine whether threads need to be serializedthrough lock-protected critical sections, and to perform thatserialization only when required. This allows the processor to exposeand exploit concurrency that is hidden in an application because ofdynamically unnecessary synchronization.

With Intel TSX, programmer-specified code regions (also referred to as“transactional regions” or just “transactions”) are executedtransactionally. If the transactional execution completes successfully,then all memory operations performed within the transactional regionwill appear to have occurred instantaneously when viewed from otherprocessors. A processor makes the memory operations of the executedtransaction, performed within the transactional region, visible to otherprocessors only when a successful commit occurs, i.e., when thetransaction successfully completes execution. This process is oftenreferred to as an atomic commit.

Since a successful transactional execution ensures an atomic commit, theprocessor executes the code region optimistically without explicitsynchronization. If synchronization was unnecessary for that specificexecution, execution can commit without any cross-thread serialization.If the processor cannot commit atomically, then the optimistic executionfails. When this happens, the processor will roll back the execution, aprocess referred to as a transactional abort. On a transactional abort,the processor will discard all updates performed in the memory regionused by the transaction, restore architectural state to appear as if theoptimistic execution never occurred, and resume executionnon-transactionally.

A processor can perform a transactional abort for numerous reasons. Aprimary reason to abort a transaction is due to conflicting memoryaccesses between the transactionally executing logical processor andanother logical processor. Such conflicting memory accesses may preventa successful transactional execution. Memory addresses read from withina transactional region constitute the read-set of the transactionalregion and addresses written to within the transactional regionconstitute the write-set of the transactional region. Intel TSXmaintains the read- and write-sets at the granularity of a cache line. Aconflicting memory access occurs if another logical processor eitherreads a location that is part of the transactional region's write-set orwrites a location that is a part of either the read- or write-set of thetransactional region. A conflicting access typically means thatserialization is required for this code region. Since Intel TSX detectsdata conflicts at the granularity of a cache line, unrelated datalocations placed in the same cache line will be detected as conflictsthat result in transactional aborts. Transactional aborts may also occurdue to limited transactional resources. For example, the amount of dataaccessed in the region may exceed an implementation-specific capacity.Additionally, some instructions and system events may causetransactional aborts. Frequent transactional aborts result in wastedcycles and increased inefficiency.

Hardware Lock Elision

Hardware Lock Elision (HLE) provides a legacy compatible instruction setinterface for programmers to use transactional execution. HLE providestwo new instruction prefix hints: XACQUIRE and XRELEASE.

With HLE, a programmer adds the XACQUIRE prefix to the front of theinstruction that is used to acquire the lock that is protecting thecritical section. The processor treats the prefix as a hint to elide thewrite associated with the lock acquire operation. Even though the lockacquire has an associated write operation to the lock, the processordoes not add the address of the lock to the transactional region'swrite-set nor does it issue any write requests to the lock. Instead, theaddress of the lock is added to the read-set. The logical processorenters transactional execution. If the lock was available before theXACQUIRE prefixed instruction, then all other processors will continueto see the lock as available afterwards. Since the transactionallyexecuting logical processor neither added the address of the lock to itswrite-set nor performed externally visible write operations to the lock,other logical processors can read the lock without causing a dataconflict. This allows other logical processors to also enter andconcurrently execute the critical section protected by the lock. Theprocessor automatically detects any data conflicts that occur during thetransactional execution and will perform a transactional abort ifnecessary.

Even though the eliding processor did not perform any external writeoperations to the lock, the hardware ensures program order of operationson the lock. If the eliding processor itself reads the value of the lockin the critical section, it will appear as if the processor had acquiredthe lock, i.e. the read will return the non-elided value. This behaviorallows an HLE execution to be functionally equivalent to an executionwithout the HLE prefixes.

An XRELEASE prefix can be added in front of an instruction that is usedto release the lock protecting a critical section. Releasing the lockinvolves a write to the lock. If the instruction is to restore the valueof the lock to the value the lock had prior to the XACQUIRE prefixedlock acquire operation on the same lock, then the processor elides theexternal write request associated with the release of the lock and doesnot add the address of the lock to the write-set. The processor thenattempts to commit the transactional execution.

With HLE, if multiple threads execute critical sections protected by thesame lock but they do not perform any conflicting operations on eachother's data, then the threads can execute concurrently and withoutserialization. Even though the software uses lock acquisition operationson a common lock, the hardware recognizes this, elides the lock, andexecutes the critical sections on the two threads without requiring anycommunication through the lock—if such communication was dynamicallyunnecessary.

If the processor is unable to execute the region transactionally, thenthe processor will execute the region non-transactionally and withoutelision. HLE enabled software has the same forward progress guaranteesas the underlying non-HLE lock-based execution. For successful HLEexecution, the lock and the critical section code must follow certainguidelines. These guidelines only affect performance; and failure tofollow these guidelines will not result in a functional failure.Hardware without HLE support will ignore the XACQUIRE and XRELEASEprefix hints and will not perform any elision since these prefixescorrespond to the REPNE/REPE IA-32 prefixes which are ignored on theinstructions where XACQUIRE and XRELEASE are valid. Importantly, HLE iscompatible with the existing lock-based programming model. Improper useof hints will not cause functional bugs though it may expose latent bugsalready in the code.

Restricted Transactional Memory (RTM) provides a flexible softwareinterface for transactional execution. RTM provides three newinstructions—XBEGIN, XEND, and XABORT—for programmers to start, commit,and abort a transactional execution.

The programmer uses the XBEGIN instruction to specify the start of atransactional code region and the XEND instruction to specify the end ofthe transactional code region. If the RTM region could not besuccessfully executed transactionally, then the XBEGIN instruction takesan operand that provides a relative offset to the fallback instructionaddress.

A processor may abort RTM transactional execution for many reasons. Inmany instances, the hardware automatically detects transactional abortconditions and restarts execution from the fallback instruction addresswith the architectural state corresponding to that present at the startof the XBEGIN instruction and the EAX register updated to describe theabort status.

The XABORT instruction allows programmers to abort the execution of anRTM region explicitly. The XABORT instruction takes an 8-bit immediateargument that is loaded into the EAX register and will thus be availableto software following an RTM abort. RTM instructions do not have anydata memory location associated with them. While the hardware providesno guarantees as to whether an RTM region will ever successfully committransactionally, most transactions that follow the recommendedguidelines are expected to successfully commit transactionally. However,programmers must always provide an alternative code sequence in thefallback path to guarantee forward progress. This may be as simple asacquiring a lock and executing the specified code regionnon-transactionally. Further, a transaction that always aborts on agiven implementation may complete transactionally on a futureimplementation. Therefore, programmers must ensure the code paths forthe transactional region and the alternative code sequence arefunctionally tested.

Detection of HLE Support

A processor supports HLE execution if CPUID.07H.EBX.HLE [bit 4]=1.However, an application can use the HLE prefixes (XACQUIRE and XRELEASE)without checking whether the processor supports HLE. Processors withoutHLE support ignore these prefixes and will execute the code withoutentering transactional execution.

Detection of RTM Support

A processor supports RTM execution if CPUID.07H.EBX.RTM [bit 11]=1. Anapplication must check if the processor supports RTM before it uses theRTM instructions (XBEGIN, XEND, XABORT). These instructions willgenerate a #UD exception when used on a processor that does not supportRTM.

Detection of XTEST Instruction

A processor supports the XTEST instruction if it supports either HLE orRTM. An application must check either of these feature flags beforeusing the XTEST instruction. This instruction will generate a #UDexception when used on a processor that does not support either HLE orRTM.

Querying Transactional Execution Status

The XTEST instruction can be used to determine the transactional statusof a transactional region specified by HLE or RTM. Note, while the HLEprefixes are ignored on processors that do not support HLE, the XTESTinstruction will generate a #UD exception when used on processors thatdo not support either HLE or RTM.

Requirements for HLE Locks

For HLE execution to successfully commit transactionally, the lock mustsatisfy certain properties and access to the lock must follow certainguidelines.

An XRELEASE prefixed instruction must restore the value of the elidedlock to the value it had before the lock acquisition. This allowshardware to safely elide locks by not adding them to the write-set. Thedata size and data address of the lock release (XRELEASE prefixed)instruction must match that of the lock acquire (XACQUIRE prefixed) andthe lock must not cross a cache line boundary.

Software should not write to the elided lock inside a transactional HLEregion with any instruction other than an XRELEASE prefixed instruction,otherwise such a write may cause a transactional abort. In addition,recursive locks (where a thread acquires the same lock multiple timeswithout first releasing the lock) may also cause a transactional abort.Note that software can observe the result of the elided lock acquireinside the critical section. Such a read operation will return the valueof the write to the lock.

The processor automatically detects violations to these guidelines, andsafely transitions to a non-transactional execution without elision.Since Intel TSX detects conflicts at the granularity of a cache line,writes to data collocated on the same cache line as the elided lock maybe detected as data conflicts by other logical processors eliding thesame lock.

Transactional Nesting

Both HLE and RTM support nested transactional regions. However, atransactional abort restores state to the operation that startedtransactional execution: either the outermost XACQUIRE prefixed HLEeligible instruction or the outermost XBEGIN instruction. The processortreats all nested transactions as one transaction.

HLE Nesting and Elision

Programmers can nest HLE regions up to an implementation specific depthof MAX_HLE_NEST_COUNT. Each logical processor tracks the nesting countinternally but this count is not available to software. An XACQUIREprefixed HLE-eligible instruction increments the nesting count, and anXRELEASE prefixed HLE-eligible instruction decrements it. The logicalprocessor enters transactional execution when the nesting count goesfrom zero to one. The logical processor attempts to commit only when thenesting count becomes zero. A transactional abort may occur if thenesting count exceeds MAX_HLE_NEST_COUNT.

In addition to supporting nested HLE regions, the processor can alsoelide multiple nested locks. The processor tracks a lock for elisionbeginning with the XACQUIRE prefixed HLE eligible instruction for thatlock and ending with the XRELEASE prefixed HLE eligible instruction forthat same lock. The processor can, at any one time, track up to aMAX_HLE_ELIDED_LOCKS number of locks. For example, if the implementationsupports a MAX_HLE_ELIDED_LOCKS value of two and if the programmer neststhree HLE identified critical sections (by performing XACQUIRE prefixedHLE eligible instructions on three distinct locks without performing anintervening XRELEASE prefixed HLE eligible instruction on any one of thelocks), then the first two locks will be elided, but the third won't beelided (but will be added to the transaction's write set). However, theexecution will still continue transactionally. Once an XRELEASE for oneof the two elided locks is encountered, a subsequent lock acquiredthrough the XACQUIRE prefixed HLE eligible instruction will be elided.

The processor attempts to commit the HLE execution when all elidedXACQUIRE and XRELEASE pairs have been matched, the nesting count goes tozero, and the locks have satisfied requirements. If execution cannotcommit atomically, then execution transitions to a non-transactionalexecution without elision as if the first instruction did not have anXACQUIRE prefix.

RTM Nesting

Programmers can nest RTM regions up to an implementation specificMAX_RTM_NEST_COUNT. The logical processor tracks the nesting countinternally but this count is not available to software. An XBEGINinstruction increments the nesting count, and an XEND instructiondecrements the nesting count. The logical processor attempts to commitonly if the nesting count becomes zero. A transactional abort occurs ifthe nesting count exceeds MAX_RTM_NEST_COUNT.

Nesting HLE and RTM

HLE and RTM provide two alternative software interfaces to a commontransactional execution capability. Transactional processing behavior isimplementation specific when HLE and RTM are nested together, e.g., HLEis inside RTM or RTM is inside HLE. However, in all cases, theimplementation will maintain HLE and RTM semantics. An implementationmay choose to ignore HLE hints when used inside RTM regions, and maycause a transactional abort when RTM instructions are used inside HLEregions. In the latter case, the transition from transactional tonon-transactional execution occurs seamlessly since the processor willre-execute the HLE region without actually doing elision, and thenexecute the RTM instructions.

Abort Status Definition

RTM uses the EAX register to communicate abort status to software.Following an RTM abort the EAX register has the following definition.

TABLE 1 RTM Abort Status Definition EAX Register Bit Position Meaning 0Set if abort caused by XABORT instruction 1 If set, the transaction maysucceed on retry, this bit is always clear if bit 0 is set 2 Set ifanother logical processor conflicted with a memory address that was partof the transaction that aborted 3 Set if an internal buffer overflowed 4Set if a debug breakpoint was hit 5 Set if an abort occurred duringexecution of a nested transaction 23:6  Reserved 31:24 XABORT argument(only valid if bit 0 set, otherwise reserved)

The EAX abort status for RTM only provides causes for aborts. It doesnot by itself encode whether an abort or commit occurred for the RTMregion. The value of EAX can be 0 following an RTM abort. For example, aCPUID instruction when used inside an RTM region causes a transactionalabort and may not satisfy the requirements for setting any of the EAXbits. This may result in an EAX value of 0.

RTM Memory Ordering

A successful RTM commit causes all memory operations in the RTM regionto appear to execute atomically. A successfully committed RTM regionconsisting of an XBEGIN followed by an XEND, even with no memoryoperations in the RTM region, has the same ordering semantics as a LOCKprefixed instruction.

The XBEGIN instruction does not have fencing semantics. However, if anRTM execution aborts, then all memory updates from within the RTM regionare discarded and are not made visible to any other logical processor.

RTM-Enabled Debugger Support

By default, any debug exception inside an RTM region will cause atransactional abort and will redirect control flow to the fallbackinstruction address with architectural state recovered and bit 4 in EAXset. However, to allow software debuggers to intercept execution ondebug exceptions, the RTM architecture provides additional capability.

If bit 11 of DR7 and bit 15 of the IA32_DEBUGCTL_MSR are both 1, any RTMabort due to a debug exception (#DB) or breakpoint exception (#BP)causes execution to roll back and restart from the XBEGIN instructioninstead of the fallback address. In this scenario, the EAX register willalso be restored back to the point of the XBEGIN instruction.

Programming Considerations

Typical programmer-identified regions are expected to transactionallyexecute and commit successfully. However, Intel TSX does not provide anysuch guarantee. A transactional execution may abort for many reasons. Totake full advantage of the transactional capabilities, programmersshould follow certain guidelines to increase the probability of theirtransactional execution committing successfully.

This section discusses various events that may cause transactionalaborts. The architecture ensures that updates performed within atransaction that subsequently aborts execution will never becomevisible. Only committed transactional executions initiate an update tothe architectural state. Transactional aborts never cause functionalfailures and only affect performance.

Instruction Based Considerations

Programmers can use any instruction safely inside a transaction (HLE orRTM) and can use transactions at any privilege level. However, someinstructions will always abort the transactional execution and causeexecution to seamlessly and safely transition to a non-transactionalpath.

Intel TSX allows for most common instructions to be used insidetransactions without causing aborts. The following operations inside atransaction do not typically cause an abort:

-   -   Operations on the instruction pointer register, general purpose        registers (GPRs) and the status flags (CF, OF, SF, PF, AF, and        ZF); and    -   Operations on XMM and YMM registers and the MXCSR register.

However, programmers must be careful when inter-mixing SSE and AVXoperations inside a transactional region. Intermixing SSE instructionsaccessing XMM registers and AVX instructions accessing YMM registers maycause transactions to abort. Programmers may use REP/REPNE prefixedstring operations inside transactions. However, long strings may causeaborts. Further, the use of CLD and STD instructions may cause aborts ifthey change the value of the DF flag. However, if DF is 1, the STDinstruction will not cause an abort. Similarly, if DF is 0, then the CLDinstruction will not cause an abort.

Instructions not enumerated here as causing abort when used inside atransaction will typically not cause a transaction to abort (examplesinclude but are not limited to MFENCE, LFENCE, SFENCE, RDTSC, RDTSCP,etc.).

The following instructions will abort transactional execution on anyimplementation:

XABORT

CPUID

PAUSE

In addition, in some implementations, the following instructions mayalways cause transactional aborts. These instructions are not expectedto be commonly used inside typical transactional regions. However,programmers must not rely on these instructions to force a transactionalabort, since whether they cause transactional aborts is implementationdependent.

-   -   Operations on X87 and MMX architecture state. This includes all        MMX and X87 instructions, including the FXRSTOR and FXSAVE        instructions.    -   Update to non-status portion of EFLAGS: CLI, STI, POPFD, POPFQ,        CLTS.    -   Instructions that update segment registers, debug registers        and/or control registers:    -   MOV to DS/ES/FS/GS/SS, POP DS/ES/FS/GS/SS, LDS, LES, LFS, LGS,        LSS, SWAPGS, WRFSBASE, WRGSBASE, LGDT, SGDT, LIDT, SIDT, LLDT,        SLDT, LTR, STR, Far CALL, Far JMP, Far RET, IRET, MOV to DRx,        MOV to CR0/CR2/CR3/CR4/CR8 and LMSW.    -   Ring transitions: SYSENTER, SYSCALL, SYSEXIT, and SYSRET.    -   TLB and Cacheability control: CLFLUSH, INVD, WBINVD, INVLPG,        INVPCID, and memory instructions with a non-temporal hint        (MOVNTDQA, MOVNTDQ, MOVNTI, MOVNTPD, MOVNTPS, and MOVNTQ).    -   Processor state save: XSAVE, XSAVEOPT, and XRSTOR.    -   Interrupts: INTn, INTO.    -   IO: IN, INS, REP INS, OUT, OUTS, REP OUTS and their variants.    -   VMX: VMPTRLD, VMPTRST, VMCLEAR, VMREAD, VMWRITE, VMCALL,        VMLAUNCH, VMRESUME, VMXOFF, VMXON, INVEPT, and INVVPID.    -   SMX: GETSEC.    -   UD2, RSM, RDMSR, WRMSR, HLT, MONITOR, MWAIT, XSETBV, VZEROUPPER,        MASKMOVQ, and V/MASKMOVDQU.        Runtime Considerations

In addition to the instruction-based considerations, runtime events maycause transactional execution to abort. These may be due to data accesspatterns or micro-architectural implementation features. The followinglist is not a comprehensive discussion of all abort causes.

Any fault or trap in a transaction that must be exposed to software willbe suppressed. Transactional execution will abort and execution willtransition to a non-transactional execution, as if the fault or trap hadnever occurred. If an exception is not masked, then that un-maskedexception will result in a transactional abort and the state will appearas if the exception had never occurred.

Synchronous exception events (#DE, #OF, #NP, #SS, #GP, #BR, #UD, #AC,#XF, #PF, #NM, #TS, #MF, #DB, #BP/INT3) that occur during transactionalexecution may cause an execution not to commit transactionally, andrequire a non-transactional execution. These events are suppressed as ifthey had never occurred. With HLE, since the non-transactional code pathis identical to the transactional code path, these events will typicallyre-appear when the instruction that caused the exception is re-executednon-transactionally, causing the associated synchronous events to bedelivered appropriately in the non-transactional execution. Asynchronousevents (NMI, SMI, INTR, IPI, PMI, etc.) occurring during transactionalexecution may cause the transactional execution to abort and transitionto a non-transactional execution. The asynchronous events will be pendedand handled after the transactional abort is processed.

Transactions only support write-back cacheable memory type operations. Atransaction may always abort if the transaction includes operations onany other memory type. This includes instruction fetches to UC memorytype.

Memory accesses within a transactional region may require the processorto set the Accessed and Dirty flags of the referenced page table entry.The behavior of how the processor handles this is implementationspecific. Some implementations may allow the updates to these flags tobecome externally visible even if the transactional region subsequentlyaborts. Some Intel TSX implementations may choose to abort thetransactional execution if these flags need to be updated. Further, aprocessor's page-table walk may generate accesses to its owntransactionally written but uncommitted state. Some Intel TSXimplementations may choose to abort the execution of a transactionalregion in such situations. Regardless, the architecture ensures that, ifthe transactional region aborts, then the transactionally written statewill not be made architecturally visible through the behavior ofstructures such as TLBs.

Executing self-modifying code transactionally may also causetransactional aborts. Programmers must continue to follow the Intelrecommended guidelines for writing self-modifying and cross-modifyingcode even when employing HLE and RTM. While an implementation of RTM andHLE will typically provide sufficient resources for executing commontransactional regions, implementation constraints and excessive sizesfor transactional regions may cause a transactional execution to abortand transition to a non-transactional execution. The architectureprovides no guarantee of the amount of resources available to dotransactional execution and does not guarantee that a transactionalexecution will ever succeed.

Conflicting requests to a cache line accessed within a transactionalregion may prevent the transaction from executing successfully. Forexample, if logical processor P0 reads line A in a transactional regionand another logical processor P1 writes line A (either inside or outsidea transactional region) then logical processor P0 may abort if logicalprocessor P1's write interferes with processor P0's ability to executetransactionally.

Similarly, if P0 writes line A in a transactional region and P1 reads orwrites line A (either inside or outside a transactional region), then P0may abort if P1's access to line A interferes with P0's ability toexecute transactionally. In addition, other coherence traffic may attimes appear as conflicting requests and may cause aborts. While thesefalse conflicts may happen, they are expected to be uncommon. Theconflict resolution policy to determine whether P0 or P1 aborts in theabove scenarios is implementation specific.

Generic Transaction Execution Embodiments:

According to “ARCHITECTURES FOR TRANSACTIONAL MEMORY”, a dissertationsubmitted to the Department of Computer Science and the Committee onGraduate Studies of Stanford University in partial fulfillment of therequirements for the Degree of Doctor of Philosophy, by Austen McDonald,June 2009, incorporated by reference herein in its entirety,fundamentally, there are three mechanisms needed to implement an atomicand isolated transactional region: versioning, conflict detection, andcontention management.

To make a transactional code region appear atomic, all the modificationsperformed by that transactional code region must be stored and keptisolated from other transactions until commit time. The system does thisby implementing a versioning policy. Two versioning paradigms exist:eager and lazy. An eager versioning system stores newly generatedtransactional values in place and stores previous memory values on theside, in what is called an undo-log. A lazy versioning system stores newvalues temporarily in what is called a write buffer, copying them tomemory only on commit. In either system, the cache is used to optimizestorage of new versions.

To ensure that transactions appear to be performed atomically, conflictsmust be detected and resolved. The two systems, i.e., the eager and lazyversioning systems, detect conflicts by implementing a conflictdetection policy, either optimistic or pessimistic. An optimistic systemexecutes transactions in parallel, checking for conflicts only when atransaction commits. A pessimistic system checks for conflicts at eachload and store. Similar to versioning, conflict detection also uses thecache, marking each line as either part of the read-set, part of thewrite-set, or both. The two systems resolve conflicts by implementing acontention management policy. Many contention management policies exist,some are more appropriate for optimistic conflict detection and some aremore appropriate for pessimistic. Described below are some examplepolicies.

Since each transactional memory (TM) system needs both versioningdetection and conflict detection, these options give rise to fourdistinct TM designs: Eager-Pessimistic (EP), Eager-Optimistic (EO),Lazy-Pessimistic (LP), and Lazy-Optimistic (LO). Table 2 brieflydescribes all four distinct TM designs.

FIGS. 11 and 12 depict an example of a multicore TM environment. 11shows many TM-enabled CPUs (CPU₁ (1114 a), CPU₂ (1114 b), etc.) on onedie (1100), connected with an interconnect (1122), under management ofan interconnect control (1120 a), (1120 b). Each CPU (1114 a), (1114 b)(also known as a Processor) may have a split cache consisting of anInstruction Cache (1116 a), (1116 b) for caching instructions frommemory to be executed and a Data Cache (1118 a), (1118 b) with TMsupport for caching data (operands) of memory locations to be operatedon by CPU (1114 a), (1114 b) (in FIG. 11, each CPU (1114 a), (1114 b)and its associated caches are referenced as (1112 a), (1112 b)). In animplementation, caches of multiple dies (1100) are interconnected tosupport cache coherency between the caches of the multiple dies (1100).In an implementation, a single cache, rather than the split cache isemployed holding both instructions and data. In implementations, the CPUcaches are one level of caching in a hierarchical cache structure. Forexample each die (1100) may employ a shared cache (1124) to be sharedamongst all the CPUs on the die (1100). In another implementation, eachdie may have access to a shared cache (1124), shared amongst all theprocessors of all the dies (1100).

FIG. 12 shows the details of an example transactional CPU environment(1212), having a CPU (1214), including additions to support TM. Thetransactional CPU (processor) (1214) may include hardware for supportingRegister Checkpoints (1226) and special TM Registers (1228). Thetransactional CPU cache may have the MESI bits (1230), Tags (1240) andData (1242) of a conventional cache but also, for example, R bits (1232)showing a line has been read by the CPU (1214) while executing atransaction and W bits (1238) showing a line has been written-to by theCPU (1214) while executing a transaction.

A key detail for programmers in any TM system is how non-transactionalaccesses interact with transactions. By design, transactional accessesare screened from each other using the mechanisms above. However, theinteraction between a regular, non-transactional load with a transactioncontaining a new value for that address must still be considered. Inaddition, the interaction between a non-transactional store with atransaction that has read that address must also be explored. These areissues of the database concept isolation.

A TM system is said to implement strong isolation, sometimes calledstrong atomicity, when every non-transactional load and store acts likean atomic transaction. Therefore, non-transactional loads cannot seeuncommitted data and non-transactional stores cause atomicity violationsin any transactions that have read that address. A system where this isnot the case is said to implement weak isolation, sometimes called weakatomicity.

Strong isolation is often more desirable than weak isolation due to therelative ease of conceptualization and implementation of strongisolation. Additionally, if a programmer has forgotten to surround someshared memory references with transactions, causing bugs, then withstrong isolation, the programmer will often detect that oversight usinga simple debug interface because the programmer will see anon-transactional region causing atomicity violations. Also, programswritten in one model may work differently on another model.

Further, strong isolation is often easier to support in hardware TM thanweak isolation. With strong isolation, since the coherence protocolalready manages load and store communication between processors,transactions can detect non-transactional loads and stores and actappropriately. To implement strong isolation in software TransactionalMemory (TM), non-transactional code must be modified to include read-and write-barriers; potentially crippling performance. Although greateffort has been expended to remove many un-needed barriers, suchtechniques are often complex and performance is typically far lower thanthat of HTMs.

TABLE 2 Transactional Memory Design Space VERSIONING Lazy Eager CONFLICTOptimistic Storing updates Not practical: DETECTION in a write buffer;waiting to update detecting conflicts memory until commit at committime. time but detecting conflicts at access time guarantees wasted workand provides no advantage Pessimistic Storing updates Updating memory,in a write buffer; keeping old detecting conflicts values in undo ataccess time. log; detecting conflicts at access time.

Table 2 illustrates the fundamental design space of transactional memory(versioning and conflict detection).

Eager-Pessimistic (EP) This first TM design described below is known asEager-Pessimistic. An EP system stores its write-set “in place” (hencethe name “eager”) and, to support rollback, stores the old values ofoverwritten lines in an “undo log”. Processors use the W (1238) and R(1232) cache bits to track read and write-sets and detect conflicts whenreceiving snooped load requests. Perhaps the most notable examples of EPsystems in known literature are LogTM and UTM.

Beginning a transaction in an EP system is much like beginning atransaction in other systems: tm_begin( ) takes a register checkpoint,and initializes any status registers. An EP system also requiresinitializing the undo log, the details of which are dependent on the logformat, but often involve initializing a log base pointer to a region ofpre-allocated, thread-private memory, and clearing a log boundsregister.

Versioning: In EP, due to the way eager versioning is designed tofunction, the MESI (1230) state transitions (cache line indicatorscorresponding to Modified, Exclusive, Shared, and Invalid code states)are left mostly unchanged. Outside of a transaction, the MESI (1230)state transitions are left completely unchanged. When reading a lineinside a transaction, the standard coherence transitions apply (S(Shared)→S, I (Invalid)→S, or I→E (Exclusive)), issuing a load miss asneeded, but the R (1232) bit is also set. Likewise, writing a lineapplies the standard transitions (S→M, E→I, I→M), issuing a miss asneeded, but also sets the W (1038) (Written) bit. The first time a lineis written, the old version of the entire line is loaded then written tothe undo log to preserve it in case the current transaction aborts. Thenewly written data is then stored “in-place,” over the old data.

Conflict Detection: Pessimistic conflict detection uses coherencemessages exchanged on misses, or upgrades, to look for conflicts betweentransactions. When a read miss occurs within a transaction, otherprocessors receive a load request; but they ignore the request if theydo not have the needed line. If the other processors have the neededline non-speculatively or have the line R (1232) (Read), they downgradethat line to S, and in certain cases issue a cache-to-cache transfer ifthey have the line in MESI's (1230) M or E state. However, if the cachehas the line W (1238), then a conflict is detected between the twotransactions and additional action(s) must be taken.

Similarly, when a transaction seeks to upgrade a line from shared tomodified (on a first write), the transaction issues an exclusive loadrequest, which is also used to detect conflicts. If a receiving cachehas the line non-speculatively, then the line is invalidated, and incertain cases a cache-to-cache transfer (M or E states) is issued. But,if the line is R (1232) or W (1038), a conflict is detected.

Validation: Because conflict detection is performed on every load, atransaction always has exclusive access to its own write-set. Therefore,validation does not require any additional work.

Commit: Since eager versioning stores the new version of data items inplace, the commit process simply clears the W (1238) and R (1232) bitsand discards the undo log.

Abort: When a transaction rolls back, the original version of each cacheline in the undo log must be restored, a process called “unrolling” or“applying” the log. This is done during tm_discard ( ) and must beatomic with regard to other transactions. Specifically, the write-setmust still be used to detect conflicts: this transaction has the onlycorrect version of lines in its undo log, and requesting transactionsmust wait for the correct version to be restored from that log. Such alog can be applied using a hardware state machine or software aborthandler.

Eager-Pessimistic has the characteristics of: Commit is simple and sinceit is in-place, very fast. Similarly, validation is a no-op. Pessimisticconflict detection detects conflicts early, thereby reducing the numberof “doomed” transactions. For example, if two transactions are involvedin a Write-After-Read dependency, then that dependency is detectedimmediately in pessimistic conflict detection. However, in optimisticconflict detection such conflicts are not detected until the writercommits.

Eager-Pessimistic also has the characteristics of: As described above,the first time a cache line is written, the old value must be written tothe log, incurring extra cache accesses. Aborts are expensive as theyrequire undoing the log. For each cache line in the log, a load must beissued, perhaps going as far as main memory before continuing to thenext line. Pessimistic conflict detection also prevents certainserializable schedules from existing.

Additionally, because conflicts are handled as they occur, there is apotential for livelock and careful contention management mechanisms mustbe employed to guarantee forward progress.

Lazy-Optimistic (LO)

Another popular TM design is Lazy-Optimistic (LO), which stores itswrite-set in a “write buffer” or “redo log” and detects conflicts atcommit time (still using the R (1032) and W (1038) bits).

Versioning: Just as in the EP system, the MESI protocol of the LO designis enforced outside of the transactions. Once inside a transaction,reading a line incurs the standard MESI transitions but also sets the R(1232) bit. Likewise, writing a line sets the W (1238) bit of the line,but handling the MESI transitions of the LO design is different fromthat of the EP design. First, with lazy versioning, the new versions ofwritten data are stored in the cache hierarchy until commit while othertransactions have access to old versions available in memory or othercaches. To make available the old versions, dirty lines (M lines) mustbe evicted when first written by a transaction. Second, no upgrademisses are needed because of the optimistic conflict detection feature:if a transaction has a line in the S state, it can simply write to itand upgrade that line to an M state without communicating the changeswith other transactions because conflict detection is done at committime.

Conflict Detection and Validation: To validate a transaction and detectconflicts, LO communicates the addresses of speculatively modified linesto other transactions only when it is preparing to commit. Onvalidation, the processor sends one, potentially large, network packetcontaining all the addresses in the write-set. Data is not sent, butleft in the cache of the committer and marked dirty (M). To build thispacket without searching the cache for lines marked W, a simple bitvector is used, called a “store buffer,” with one bit per cache line totrack these speculatively modified lines. Other transactions use thisaddress packet to detect conflicts: if an address is found in the cacheand the R (1232) and/or W (1238) bits are set, then a conflict isinitiated. If the line is found but neither R (1232) nor W (1238) isset, then the line is simply invalidated, which is similar to processingan exclusive load.

To support transaction atomicity, these address packets must be handledatomically, i.e., no two address packets may exist at once with the sameaddresses. In an LO system, this can be achieved by simply acquiring aglobal commit token before sending the address packet. However, atwo-phase commit scheme could be employed by first sending out theaddress packet, collecting responses, enforcing an ordering protocol(perhaps oldest transaction first), and committing once all responsesare satisfactory.

Commit: Once validation has occurred, commit needs no special treatment:simply clear W (1238) and R (1232) bits and the store buffer. Thetransaction's writes are already marked dirty in the cache and othercaches' copies of these lines have been invalidated via the addresspacket. Other processors can then access the committed data through theregular coherence protocol.

Abort: Rollback is equally easy: because the write-set is containedwithin the local caches, these lines can be invalidated, then clear W(1238) and R (1232) bits and the store buffer. The store buffer allows Wlines to be found to invalidate without the need to search the cache.

Lazy-Optimistic has the characteristics of: Aborts are very fast,requiring no additional loads or stores and making only local changes.More serializable schedules can exist than found in EP, which allows anLO system to more aggressively speculate that transactions areindependent, which can yield higher performance. Finally, the latedetection of conflicts can increase the likelihood of forward progress.

Lazy-Optimistic also has the characteristics of: Validation takes globalcommunication time proportional to size of write set. Doomedtransactions can waste work since conflicts are detected only at committime.

Lazy-Pessimistic (LP)

Lazy-Pessimistic (LP) represents a third TM design option, sittingsomewhere between EP and LO: storing newly written lines in a writebuffer but detecting conflicts on a per access basis.

Versioning: Versioning is similar but not identical to that of LO:reading a line sets its R bit (1032), writing a line sets its W bit(1238), and a store buffer is used to track W lines in the cache. Also,dirty (M) lines must be evicted when first written by a transaction,just as in LO. However, since conflict detection is pessimistic, loadexclusives must be performed when upgrading a transactional line from I,S→M, which is unlike LO.

Conflict Detection: LP's conflict detection operates the same as EP's:using coherence messages to look for conflicts between transactions.

Validation: Like in EP, pessimistic conflict detection ensures that atany point, a running transaction has no conflicts with any other runningtransaction, so validation is a no-op.

Commit: Commit needs no special treatment: simply clear W (1238) and R(1232) bits and the store buffer, like in LO.

Abort: Rollback is also like that of LO: simply invalidate the write-setusing the store buffer and clear the W and R bits and the store buffer.

Eager-Optimistic (EO)

The LP has the characteristics of: Like LO, aborts are very fast. LikeEP, the use of pessimistic conflict detection reduces the number of“doomed” transactions. Like EP, some serializable schedules are notallowed and conflict detection must be performed on each cache miss.

The final combination of versioning and conflict detection isEager-Optimistic (EO). EO may be a less than optimal choice for HTMsystems: since new transactional versions are written in-place, othertransactions have no choice but to notice conflicts as they occur (i.e.,as cache misses occur). But since EO waits until commit time to detectconflicts, those transactions become “zombies,” continuing to execute,wasting resources, yet are “doomed” to abort.

EO has proven to be useful in STMs and is implemented by Bartok-STM andMcRT. A lazy versioning STM needs to check its write buffer on each readto ensure that it is reading the most recent value. Since the writebuffer is not a hardware structure, this is expensive, hence thepreference for write-in-place eager versioning. Additionally, sincechecking for conflicts is also expensive in an STM, optimistic conflictdetection offers the advantage of performing this operation in bulk.

Contention Management

How a transaction rolls back once the system has decided to abort thattransaction has been described above, but, since a conflict involves twotransactions, the topics of which transaction should abort, how thatabort should be initiated, and when should the aborted transaction beretried need to be explored. These are topics that are addressed byContention Management (CM), a key component of transactional memory.Described below are policies regarding how the systems initiate abortsand the various established methods of managing which transactionsshould abort in a conflict.

Contention Management Policies

A Contention Management (CM) Policy is a mechanism that determines whichtransaction involved in a conflict should abort and when the abortedtransaction should be retried. For example, it is often the case thatretrying an aborted transaction immediately does not lead to the bestperformance. Conversely, employing a back-off mechanism, which delaysthe retrying of an aborted transaction, can yield better performance.STMs first grappled with finding the best contention management policiesand many of the policies outlined below were originally developed forSTMs.

CM Policies draw on a number of measures to make decisions, includingages of the transactions, size of read- and write-sets, the number ofprevious aborts, etc. The combinations of measures to make suchdecisions are endless, but certain combinations are described below,roughly in order of increasing complexity.

To establish some nomenclature, first note that in a conflict there aretwo sides: the attacker and the defender. The attacker is thetransaction requesting access to a shared memory location. Inpessimistic conflict detection, the attacker is the transaction issuingthe load or load exclusive. In optimistic, the attacker is thetransaction attempting to validate. The defender in both cases is thetransaction receiving the attacker's request.

An Aggressive CM Policy immediately and always retries either theattacker or the defender. In LO, Aggressive means that the attackeralways wins, and so Aggressive is sometimes called committer wins. Sucha policy was used for the earliest LO systems. In the case of EP,Aggressive can be either defender wins or attacker wins.

Restarting a conflicting transaction that will immediately experienceanother conflict is bound to waste work—namely interconnect bandwidthrefilling cache misses. A Polite CM Policy employs exponential backoff(but linear could also be used) before restarting conflicts. To preventstarvation, a situation where a process does not have resourcesallocated to it by the scheduler, the exponential backoff greatlyincreases the odds of transaction success after some n retries.

Another approach to conflict resolution is to randomly abort theattacker or defender (a policy called Randomized). Such a policy may becombined with a randomized backoff scheme to avoid unneeded contention.

However, making random choices, when selecting a transaction to abort,can result in aborting transactions that have completed “a lot of work”,which can waste resources. To avoid such waste, the amount of workcompleted on the transaction can be taken into account when determiningwhich transaction to abort. One measure of work could be a transaction'sage. Other methods include Oldest, Bulk TM, Size Matters, Karma, andPolka. Oldest is a simple timestamp method that aborts the youngertransaction in a conflict. Bulk TM uses this scheme. Size Matters islike Oldest but instead of transaction age, the number of read/writtenwords is used as the priority, reverting to Oldest after a fixed numberof aborts. Karma is similar, using the size of the write-set aspriority. Rollback then proceeds after backing off a fixed amount oftime. Aborted transactions keep their priorities after being aborted(hence the name Karma). Polka works like Karma but instead of backingoff a predefined amount of time, it backs off exponentially more eachtime.

Since aborting wastes work, it is logical to argue that stalling anattacker until the defender has finished their transaction would lead tobetter performance. Unfortunately, such a simple scheme easily leads todeadlock.

Deadlock avoidance techniques can be used to solve this problem. Greedyuses two rules to avoid deadlock. The first rule is, if a firsttransaction, T1, has lower priority than a second transaction, T0, or ifT1 is waiting for another transaction, then T1 aborts when conflictingwith T0. The second rule is, if T1 has higher priority than T0 and isnot waiting, then T0 waits until T1 commits, aborts, or starts waiting(in which case the first rule is applied). Greedy provides someguarantees about time bounds for executing a set of transactions. One EPdesign (LogTM) uses a CM policy similar to Greedy to achieve stallingwith conservative deadlock avoidance.

Example MESI coherency rules provide for four possible states in which acache line of a multiprocessor cache system may reside, M, E, S, and I,defined as follows:

Modified (M): The cache line is present only in the current cache, andis dirty; it has been modified from the value in main memory. The cacheis required to write the data back to main memory at some time in thefuture, before permitting any other read of the (no longer valid) mainmemory state. The write-back changes the line to the Exclusive state.

Exclusive (E): The cache line is present only in the current cache, butis clean; it matches main memory. It may be changed to the Shared stateat any time, in response to a read request. Alternatively, it may bechanged to the Modified state when writing to it.

Shared (S): Indicates that this cache line may be stored in other cachesof the machine and is “clean”; it matches the main memory. The line maybe discarded (changed to the Invalid state) at any time.

Invalid (I): Indicates that this cache line is invalid (unused).

TM coherency status indicators (R (1232), W (1238)) may be provided foreach cache line, in addition to, or encoded in the MESI coherency bits.An R (1232) indicator indicates the current transaction has read fromthe data of the cache line, and a W (1238) indicator indicates thecurrent transaction has written to the data of the cache line.

In another aspect of TM design, a system is designed using transactionalstore buffers. U.S. Pat. No. 6,349,361 titled “Methods and Apparatus forReordering and Renaming Memory References in a Multiprocessor ComputerSystem,” filed Mar. 31, 2000 and incorporated by reference herein in itsentirety, teaches a method for reordering and renaming memory referencesin a multiprocessor computer system having at least a first and a secondprocessor. The first processor has a first private cache and a firstbuffer, and the second processor has a second private cache and a secondbuffer. The method includes the steps of, for each of a plurality ofgated store requests received by the first processor to store a datum,exclusively acquiring a cache line that contains the datum by the firstprivate cache, and storing the datum in the first buffer. Upon the firstbuffer receiving a load request from the first processor to load aparticular datum, the particular datum is provided to the firstprocessor from among the data stored in the first buffer based on anin-order sequence of load and store operations. Upon the first cachereceiving a load request from the second cache for a given datum, anerror condition is indicated and a current state of at least one of theprocessors is reset to an earlier state when the load request for thegiven datum corresponds to the data stored in the first buffer.

The main implementation components of one such transactional memoryfacility are a transaction-backup register file for holdingpre-transaction GR (general register) content, a cache directory totrack the cache lines accessed during the transaction, a store cache tobuffer stores until the transaction ends, and firmware routines toperform various complex functions. In this section a detailedimplementation is described.

IBM zEnterprise EC12 Enterprise Server Embodiment

The IBM zEnterprise EC12 enterprise server introduces transactionalexecution (TX) in transactional memory, and is described in part in apaper, “Transactional Memory Architecture and Implementation for IBMSystem z” of Proceedings Pages 25-36 presented at MICRO-45, 1-5 Dec.2012, Vancouver, British Columbia, Canada, available from IEEE ComputerSociety Conference Publishing Services (CPS), which is incorporated byreference herein in its entirety.

Table 3 shows an example transaction. Transactions started with TBEGINare not assured to ever successfully complete with TEND, since they canexperience an aborting condition at every attempted execution, e.g., dueto repeating conflicts with other CPUs. This requires that the programsupport a fallback path to perform the same operationnon-transactionally, e.g., by using traditional locking schemes. Thisputs significant burden on the programming and software verificationteams, especially where the fallback path is not automatically generatedby a reliable compiler.

TABLE 3 Example Transaction Code LHI R0,0 *initialize retry count=0 loopTBEGIN *begin transaction JNZ abort *go to abort code if CC1=0 LT R1,lock *load and test the fallback lock JNZ lckbzy *branch if lock busy .. . perform operation . . . TEND *end transaction . . . . . . . . . . .. lckbzy TABORT *abort if lock busy; this *resumes after TBEGIN abort JOfallback *no retry if CC=3 AHI R0, 1  *increment retry count CIJNL R0,6,fallback  *give up after 6 attempts PPA R0, TX *random delay based onretry count . . . potentially wait for lock to become free . . . J loop *jump back to retry fallback OBTAINlock *using Compare&Swap . . .perform operation . . . RELEASE lock . . . . . . . . . . . .

The requirement of providing a fallback path for aborted TransactionExecution (TX) transactions can be onerous. Many transactions operatingon shared data structures are expected to be short, touch only a fewdistinct memory locations, and use simple instructions only. For thosetransactions, the IBM zEnterprise EC12 introduces the concept ofconstrained transactions; under normal conditions, the CPU (1214) (FIG.12) assures that constrained transactions eventually end successfully,albeit without giving a strict limit on the number of necessary retries.A constrained transaction starts with a TBEGINC instruction and endswith a regular TEND. Implementing a task as a constrained ornon-constrained transaction typically results in very comparableperformance, but constrained transactions simplify software developmentby removing the need for a fallback path. IBM's Transactional Executionarchitecture is further described in z/Architecture, Principles ofOperation, Tenth Edition, SA22-7832-09 published September 2012 fromIBM, incorporated by reference herein in its entirety.

A constrained transaction starts with the TBEGINC instruction. Atransaction initiated with TBEGINC must follow a list of programmingconstraints; otherwise the program takes a non-filterableconstraint-violation interruption. Exemplary constraints may include,but not be limited to: the transaction can execute a maximum of 32instructions, all instruction text must be within 256 consecutive bytesof memory; the transaction contains only forward-pointing relativebranches (i.e., no loops or subroutine calls); the transaction canaccess a maximum of 4 aligned octowords (an octoword is 32 bytes) ofmemory; and restriction of the instruction-set to exclude complexinstructions like decimal or floating-point operations. The constraintsare chosen such that many common operations like doubly linkedlist-insert/delete operations can be performed, including the verypowerful concept of atomic compare-and-swap targeting up to 4 alignedoctowords. At the same time, the constraints were chosen conservativelysuch that future CPU implementations can assure transaction successwithout needing to adjust the constraints, since that would otherwiselead to software incompatibility.

TBEGINC mostly behaves like XBEGIN in TSX or TBEGIN on IBM's zEC12servers, except that the floating-point register (FPR) control and theprogram interruption filtering fields do not exist and the controls areconsidered to be zero. On a transaction abort, the instruction addressis set back directly to the TBEGINC instead of to the instruction after,reflecting the immediate retry and absence of an abort path forconstrained transactions.

Nested transactions are not allowed within constrained transactions, butif a TBEGINC occurs within a non-constrained transaction it is treatedas opening a new non-constrained nesting level just like TBEGIN would.This can occur, e.g., if a non-constrained transaction calls asubroutine that uses a constrained transaction internally. Sinceinterruption filtering is implicitly off, all exceptions during aconstrained transaction lead to an interruption into the operatingsystem (OS). Eventual successful finishing of the transaction relies onthe capability of the OS to page-in the at most 4 pages touched by anyconstrained transaction. The OS must also ensure time-slices long enoughto allow the transaction to complete.

TABLE 4 Transaction Code Example TBEGINC *begin constrained transaction. . . perform operation . . . TEND  *end transaction

Table 4 shows the constrained-transactional implementation of the codein Table 3, assuming that the constrained transactions do not interactwith other locking-based code. No lock testing is shown therefore, butcould be added if constrained transactions and lock-based code weremixed.

When failure occurs repeatedly, software emulation is performed usingmillicode as part of system firmware. Advantageously, constrainedtransactions have desirable properties because of the burden removedfrom programmers.

Transactional execution uses tracked read and write sets to determine ifexecution can be guaranteed without interaction with another process. Atransaction is a sequence of instructions that look like they have allbeen executed without any intervening interaction with another process.Interference relates to memory location conflicts with another operationon the same memory location. Recovery from interference is expensive,and typically consists of aborting a transaction and restarting. At thesame time, an over-indication of interference leads to aborting atransaction where there is no actual interference. In prior arttransactional memory implementations, read sets R (1232) and write setsW (1238) are updated when in instruction executes. In accordance withthe present invention, pending address queues are used or created toaddress and remedy unnecessary interference, and specifically related tothe speculative execution of instructions within the transaction. When aspeculative instruction becomes non-speculative with respect to otherinstructions preceding it (but not relative to the possibility of atransaction abort), an associated address is transferred from a pendingaddress queue to the R or W set for a read and write memory access,respectively. In one embodiment, the address queues may be combined intoone shared queue. Similarly, in one embodiment, the address queues canbe shared with other queue structures. In one embodiment, these queuesare shared with at least of a store queue or store order queue, e.g.,for tracking W sets, and a load order queue, e.g., for tracking R sets.Non-speculative indicators are maintained in the cache to track R and Wsets corresponding to reads and writes to memory by non-speculativeinstructions (i.e., instructions which are no long speculative). Asshown and described herein, read set bits and write set bits aremaintained in the cache in association with non-speculative read andwrite memory accesses, respectively, and associated read and writeaddresses are placed in the pending address queue(s) in association withspeculative read and write memory accesses. Interference is managed andresolved by checking bits set in the cache and/or an address in thequeue depending on the context of the received request.

Referring to FIG. 1, a flow chart (100) is provided illustrating aprocess for determining a potential conflict associated with theexecution of transactions on at least two processors, a local processor(1114 a) and a remote processor (1114 b) as shown in FIG. 11, in atransactional memory system. As shown, a remote processor requests data(102). In response to receiving such a request, logic operativelycoupled to a local processor accesses at least one local cache (104). Itis then determined if the data is present in the at least one localcache (106). If the data is not present, the flow ends (108). However,if the data is present, it is then ascertained if the request receivedfrom the remote processor is a read request or a write request (110).Multiple transactions can read the same data. To determine whetherinterference exists for a read request, a set of alternative tests areconducted. First a test is conducted whether the request received instep (102) corresponds to a read or a write request. If the requestreceived at (102) corresponds to a read request, control transfers tostep (112). Alternatively, if the request received at (102) correspondsto a write, control transfers to step (120).

In order to determine interference for a received read request, at step(112) it is determined if a non-speculative write indicator, e.g., thewrite set (1238) of FIG. 12, or the TX DIRTY indicator (1052) of FIG.,has been set in the cache by a local processor. Presence of the bit isevidence that requested data is the subject of a pending memory writerequest corresponding to a non-speculative instruction (non-speculativewith respect to other instructions, but speculative with respect to atransaction abort or rollback occurring) of a transaction on the presentprocessor, indicating that a transaction on the local processor is inthe process of changing the requested data associated with thetransaction. A second determination (116) is made to determine whetherthe address corresponding to remote request received in step (102)corresponds to an address in a the pending write queue corresponding tomemory write requested executed by instructions which are speculativerelative to other instructions, and speculative relative to the possibleoccurrence of a transaction abort or transaction rollback. If either ofthe determinations at step (112) and step (116) is affirmative, recoveryof the transaction is initiated due to a conflict (114). Anon-affirmative query at step (112) is followed by the query at step(116). If the query at step (116) is not affirmative, this is anindication that there is no interference associated with the readrequest. A cache line for the requested access is provided to therequesting processor (126), and the flow for checking for a conflictbetween a request from a remote processor and a currently activetransaction on a local processor concludes (128). Accordingly, for aread request, the determination looks for a non-speculative write setindicator of non-speculative memory write access(es), e.g., as indicatedby W set (1238) or TX-DIRTY (1052), in the cache (the write accessesbeing non-speculative with respect to other instructions, butspeculative with respect to the possibility of a transaction abort ortransaction rollback occurring) and speculatively executed memory writeaccess instructions (the instructions being speculative with respect toother instructions, and to the possibility of a transaction abort orrollback) corresponding to a speculative member of a write set pendingin the write queue.

The determination for a potential conflict for a write request isdifferent from that of a read request. More specifically and as shown atstep (120) and step (124), the write request interference determinationtests for presence of interference with respect to non-speculativeinstructions of an active transaction, and speculative instructions of atransaction. At step (120), the method performs a determination forinterference of the request of (102) with either a non-speculative writeset indicator, e.g., W set (1238) or TX DIRTY (1052), or anon-speculative read set indicator e.g., R set (1232) or TX READ (1048),in the cache, corresponding to non-speculative memory read or memorywrite instructions (including but not limited to load and storeinstructions, respectively, and computational instructions having atleast a memory write and memory read, operand, respectively, thenon-speculative nature referring to the instruction beingnon-speculative with respect to other instructions, but beingspeculative with respect to the possibility of a transaction abort ortransaction rollback occurring). At step (124), which takes placefollowing a non-affirmative response to the determination at step (120),an address corresponding to the request of (102) is compared to addresspending queues corresponding to speculatively executed memory read ormemory write instructions (including but not limited to load and storeinstructions, respectively, and computational instructions having atleast a memory write and memory read, operand, respectively, thespeculative nature referring to the instruction being speculative withrespect to other instructions, and also being speculative with respectto the possibility of a transaction abort or transaction rollbackoccurring). If either of determinations (120) and (124) are affirmative,corresponding as indicated with an associated address pending in one ofthe read and write queues, respectively, an interference of a writerequest received at (102) from a remote processor with an activetransaction of the present local processor has been detected. In oneembodiment, the read and write queues may be combined into a singlequeue. The read set indicator pertains to reading new data versus olddata, in order to ensure a sequentially consistent transaction orderingwith respect to write requests from other processors. The write setindicator tests and validates that two transactions cannot write to thesame location concurrently. A positive response to at least one of theevaluation(s) at step (120) and step (124) is followed by initiating arecovery due to conflict (122). However, a negative response to bothevaluation(s) at step (120) and (124) is followed by proceeding to step(126) for completion of the transaction. Accordingly, as shown as steps(112) and (120), the transaction interference resolution for read andwrite requests evaluates at least the write queue for a pending addresscorresponding to the received request and the cache for setting of anon-speculative indicator, and performs the read or write transactionbased on clearance of the potential conflict.

As shown and described in FIG. 1, both a pending read queue (1082) and apending write queue (1084) are created and used to resolve speculativetransactions. More specifically, as shown and described in FIG. 1, thepending read and write queues are created and used to store addressescorresponding to speculative load and store instructions, respectively.The queues may be one shared queue of a pending request, or separateload and store pending queues. In one embodiment, the queue(s) can beshared with other queue structures, such as a store queue or load orderqueue. For both read and write requests, if at least one addresscorresponding to speculative member of a write set pending is presentand detected in the write queues, then a potential conflict is detectedand recovery is initiated.

Determination of an address conflict further takes into account both thesize of the remote request received at (102), and of the size of theoperation of the pending queue. In one embodiment, when transactionalread and write sets are maintained and tracked with a granularity of acache line two addresses correspond in the address determinations (116)and (124) when the request address of the request received at (102) andthe addresses stored in the pending address queues refer to the samecache line. In at least one embodiment, pending address queues aremaintained only with a number of bits to uniquely identify a cache line,and low order bits are not stored. In another embodiment, low order bitsare stored. In particular, in an embodiment where an address queueserves several purposes, e.g., as a store queue and a write addressqueue, all bits may be maintained. In yet another embodiment, some highorder bits are not stored in the pending address queue, creating therisk of vale matches. However, omitting high order bits may help improvearea and power consumption as well as latency. In another embodiment,upper bits might be hashed, and a hashed set of high order bits may bestored in a pending address queue. In one such embodiment, the receivedrequests high order bits are also hashed, and hashes are compared. Inthese embodiments, a hash is a unique code created to represent aplurality of high order bits with fewer bits while preferably maximizingdisambiguation. A variety of hash codes, such as including but notlimited to the secure hashing algorithm 1 (SHA1) are known in the art,and may be used in conjunction with the present invention. In otherembodiment, interference is tracked at a granularity other than cachelines (e.g., smaller than cache lines), and two addresses are indicatedto interfere with each when they refer to the same granularity.

Referring to FIG. 2, a flow chart (200) is provided illustrating theflow for instructions through a microprocessor while managing aspeculative access queue(s). As shown, an instruction is fetched anddecoded (202), and availability of instruction inputs is ascertained(204). An instruction is issued (206) and speculatively executed (208).The instruction may be in the form of a memory read request or a memorywrite request. In both forms, an address associated with the speculativeexecution is placed and stored in an associated queue, including theread access or write access pending queue corresponding to thespeculative load or store instruction(s) or other instructions withmemory read and write operands (210). In one embodiment, the speculativeinstruction is addressed by a tag, e.g. a load tag and/or a store tag.In one embodiment, the load and/or storage tag is the same tag as usedfor managing load and store queue(s). In another embodiment, the loadaccess and store access pending queue is shared with a load order andstore queue, or other pre-existing queues. Following placement of theaddress in the queue(s) at step (210), a result write back takes place(212), preferably using some buffering mechanism. In one embodiment, thebuffering mechanism for speculative execution includes rename registersor future file(s). Following step (212), it is determined if atransaction interference is present so as to cause a flush of thespeculative instruction (214). The instruction is then buffered, e.g.,in a global completion table in accordance with prior art, until theinstruction is ready to complete. While the instruction is waiting tocomplete, the instruction is a speculative instruction with respect toother instructions, and also speculative with respect to the possibilityof a transaction abort or rollback occurring. During this period, theinstruction is constantly compared against a variety of upset events,including branch mis-prediction and exception conditions that may affectthe validity of the instruction execution. If a branch mis-prediction,exception or other upset event is detected during this period, controlpasses to step (220).

A negative response at step (214) is an indication that the instructionis complete or next to complete and is followed by removing the addresscorresponding to the present instruction from the read or write accessqueue, also referred to as a pending queue representing speculativeaccesses (216) with respect to other instructions. In addition, step(216) includes updating the read set and write set indicator(s) in thecache for cache lines corresponding to this instruction's access (216),and committing the result (218). In one embodiment, the update of theindicator(s) in the cache is deferred until execution of theinstruction(s) is committed. When an instruction is committed, it nolonger is speculative with respect to other instructions, but remainsspeculative with respect to the possibility of a transaction abort orrollback occurring. When a transaction rolls back (not shown), stateupdates corresponding to the plurality of updates effected byinstructions within a transaction are rolled back. When a transactioncompletes (not shown), instruction effects may no longer be rolled back.The act of committing the result of the entire transaction (not shown)includes providing the associated cache line for the requested access tothe remote processor.

The determination at step (214) pertains to assessing the state of thespeculative instruction. A positive response at step (214) is followedby flushing this instruction and subsequent instructions (220), removingthe address corresponding to the present instruction and all subsequentinstructions from the read or write access queue(s) (222), andreinitiating execution (224). In addition, updates to future files,rename registers, store queues, and other processor structures havingbeen affected by the present instruction (and its subsequentinstructions) may be similarly invalidated in conjunction with theflushing of the present instruction. In a difference from step (216),the read set and write set bits are not updated. Thereafter, executionof the present instruction, or of another instruction, is reinitiated.Accordingly, execution of the speculative instruction as shown hereinplaces an address associated with the speculative execution in the readaccess or write access pending queue, and subsequently removes theaddress from the queue(s) depending on the basis of the flush of thespeculative instruction.

As shown and described in FIGS. 1 and 2, one or more pending queues areutilized to manage speculative instructions. Referring to FIG. 3, ablock diagram (300) is provided illustrating a pending access queue. Asshown herein, addresses to be entered into the pending read and writeaddress queues are received from execution units of the local processor(310) and are stored in the queue (320), which is shown with emptyaddress slots (332) and (334) and filled address slots (336), (338),(340), and (342). The quantity of empty and filled address slots are fordescriptive purposes only and should not be considered limiting. Thequeue (320) is shown in communication with cache (350), which as shownand described in FIG. 4 contains read and write set bits (e.g., (1232)and (1238), respectively) that are updated with the processing ofinstructions.

The queue(s) shown and described in FIG. 3 is employed for interferencemanagement of instructions. In one embodiment, the queue(s) is a firstin first out queue. Referring to FIG. 4, a set of flow charts (400) isprovided illustrating methods performed in conjunction with themanagement of a pending address queue in accordance with the presentinvention. Although three separate flow charts are shown, the processesdepicted by the charts take place concurrently. The method of flow chart(410) pertains to a received address. As shown, a new instruction isstarted (412), an associated load or store address for a speculativeinstruction is received (414), and the address is placed in the queue(416). Following step (416), it is determined if the received addresscorresponds to a speculative load instruction (418). An affirmativeresponse is followed by adding the received address to the tail of thequeue (420), and a non-affirmative response ends the management of thepending address queue for the received instruction (422). As shown, theprocess waits for receipt of a corresponding address and places theaddress at the end of the queue. The test shown herein for placement inthe queue addresses whether there is an address that was received.Accordingly, receipt of the address enables the address to be placed inthe queue for a pending speculative instruction.

In at least one embodiment, the method may be further augmented tohandle a queue full condition. In one embodiment, when a queue is full,the processor stalls. In another embodiment, when a queue is full, afree queue entry may be created by committing one queue entry, e.g.,such as performing step (436) on one of a front queue address, a middlequeue address, a random queue address, and the received address, therebyprematurely committing an entry to R or W sets to avoid the performancepenalty associated with a stall condition, but at the cost of increasingthe risk of interference over indication and incurring the cost ofperforming unnecessary transaction aborts. An embodiment of queueoverflow management is disclosed in detail herein below in conjunctionwith FIGS. 6 and 7.

The queue management process continually waits for notification that aspeculative instruction has completed. As shown at flow chart (430),steps are shown to describe processing a commit instruction. It isdetermined if an address at the front of the queue corresponds to aspeculative instruction being committed (432), e.g., by step (218). Anegative response is followed by returning to the query at step (432).Substantially, the instruction commit process waits for the instructioncorresponding to the queue entry at the front of the pending addressqueue entry to be committed. A positive response to the test at step(432) is followed by completing the instruction, in conjunction withupdating the read set bit or write set bit in the cache (434) andremoving the address corresponding to a completed instruction from thefront entry of the pending address queue (436). When work on theinstruction is completed, the address is and removed from the front ofthe queue. Accordingly, processing the commit instruction is indicatedby both setting an appropriately indicated bit in the cache and updatingthe queue.

For various reasons and circumstance an instruction may not complete.The flow chart (450) addresses the steps of processing the flushing ofan instruction that takes place when a started instruction will notcomplete. As shown, it is determined if a flush indication has beenreceived (452). At such time as it is received, the first entrycorresponding to the flush instruction is ascertained (454), and thisentry and all subsequent entries are removed from the queue(s) (456).Following step (456), the tail pointer is reset (458). Following thereset at step (458) or a negative response to the determination at step(452), the process waits for receipt of the next flush instruction. Theflush instruction is processed when it is determined that theinstruction will not complete, and the instruction must be restarted.

The addresses in the queue are referred to herein as pertaining tospeculative instructions when speculative instructions are speculativerelative to other instructions (as well as to the possibility of atransaction abort or rollback occurring). Referring to FIG. 5, a flowchart (500) is provided illustrating steps for managing the pendingqueue, and specifically performing interference checks with respect toaddresses corresponding to received requests from remote processors. Asshown, a communication is received to check for interference of one ormore addresses in the queue (502). It is determined if a remote requesthas been received for a particular address (504), and if the request hasbeen received, it is determined if the address associated with therequest corresponds to any addresses present in the queue (506). Apositive response to the determination at step (506) is an indication ofan interference (508). However, a negative response to either of thedeterminations indicates that there is no interference and the processmay proceed having passed the interference check (510). Accordingly, theinterference check entails reviewing the queue to determine if anaddress is present, as such presence is an indication of possibleinterference that may warrant an invalidation of the transaction. In atleast one embodiment, two addresses are determined to interfere witheach other when they correspond to the same granule of coherencemanagement, or transactional interference tracking. In one embodiment,coherence tracking and/or transaction state tracking is performed at acache line level. In another embodiment, it is performed at a sub-cacheline level. Yet other embodiments may use yet other coherence granules.

As shown and described in FIGS. 1-5, a pending address queue is providedand employed to manage interference associated with a transaction. It isunderstood that the queue itself requires management. Referring to FIG.6 a flow chart (600) is provided illustrating management of the queue,and more specifically, addressing overflow of the queue. The queue is apending address queue, and the size of the queue is finite. The queuemay be subject to overflow or configured to make forward progress andtolerate some over-indication. As shown, it is determined if an addresscorresponding to a speculative load is received in the queue (602). Anegative response to the determination ends the management query forthis transaction in relation to the speculative load (604). In oneembodiment, review of the queue for overflow is conducted in response toreceipt of a load or store (or other memory access instruction. However,a positive response to the determination at step (602) is following bydetermining if the queue is full (606). A full queue may be problematic,since interference is established based on reviewing the queue for acorresponding address. If the queue is full, the process may need tostall (608) until such time as there is space in the queue for thespeculative load, after which the process returns to step (606). Thestall may end at such time as a prior instruction address in the queue,e.g. at the front of the queue, is completed. At such time as there isspace in the queue, the received address associated with the speculativeload is added to the tail of the queue (610), since the queue is a firstin first out queue. In one embodiment, the stall may be replaced by someform of forward progress, so that there is a toleration ofover-indication. The management of the queue as shown herein is usefulwhen the queue is shared with a load order queue, a store order queue,etc. Accordingly, the queue requires some management to ensure thatreceived instructions are properly managed and not lost or otherwisediscarded in view of a full queue.

Referring to FIG. 7, a flow chart (700) is provided illustrating anotherform of management of the pending address queue that does not stall theupdate of the queue, as shown and described in FIG. 6. As shown herein,it is determined if an address corresponding to a speculative load isreceived in the queue (702). A negative response to the determinationends the management query for this transaction in relation to thespeculative load (704). More specifically, either there is notransaction or the received transaction does not correspond to aspeculative transaction that would entail use of the queue. A positiveresponse to the determination is followed by a query to determine if thequeue is full (706). As described above, a full queue requiresmanagement so that the received instructions are not lost or discarded.If the queue is full, the write set or read set bit is updated in thecache to reflect an address at the front of the queue as part of thewrite set of read set (708). After the cache is updated, the associatedentry is removed from the queue (710), i.e. the front entry of the queueis removed, thereby making space for the incoming address. The processshown herein optimizes the performance by moving the oldest speculativeaddress in the queue to write set or read set. In one embodiment, theincoming speculative address may be moved to the write set or read set.Following step (710) or an indication that the queue is not full, e.g. anegative response to the determination at step (706), the receivedaddress associated with the speculative load is added to the tail of thequeue (712). Accordingly, the process for management of the queue maytake on different forms while considering the age of the addresses inthe queue. In another aspect of the present invention, the same processis applied for a pending store address queue when a store instruction isreceived, or for a combined load/store address queue when a load or astore address is received.

As shown and described in FIGS. 1-7 instructions are received in one ormore pending address queues, and are flushed from the queue based ondifferent scenarios. However, in some circumstance, the speculativeinstruction may be in a state where it is not known whether there willbe an over-indication. Specifically, the over-indication addresses apotential transaction interference exclusively conflicting with aspeculative instruction.

Referring to FIGS. 8A-8C, a flow chart (800) is provided illustrating aprocess for holding a response to a requesting processor based onpotential interference associated with a transaction. Similar to theprocess shown and described in FIG. 1, a remote processor requests data(802). In response to receiving such a request, logic operativelycoupled to a local processor accesses at least one local cache (804). Itis then determined if the data is present in the accessed local cache(806). If the data is not present, the flow ends (808). However, if thedata is present, it is then ascertained if the request received from theremote processor requests the data for a read request or a write request(810). Multiple transactions can read the same data. To determinewhether interference exits for a read request, a set of alternativetests is conducted. First, a test is conducted to determine whether therequest received in step (802) corresponds to a read or a write request.If the request received at step (802) corresponds to a read request,control transfers to step (812). Alternatively, if the request receivedat step (802) corresponds to a write, control transfers to step (826).

In order to determine interference for a received read request, at step(812) it is determined if a non-speculative write indicator, e.g., thewrite set (1238) of FIG. 12, or the TX DIRTY indicator (1052) of FIG.10, has been set in the cache by a local processor. Presence of the bitis evidence that that requested data is the subject of a pending memorywrite request corresponding to a non-speculative instruction(non-speculative with respect to other instructions, but speculativewith respect to a transaction abort or rollback occurring) of atransaction on the present processor, indicating that a transaction onthe local processor may be in the process of changing the requested dataassociated with the transaction. An affirmative response to thedetermination at step (812) is followed by initiating recovery of thetransaction due to a conflict (814). If the response at step (812) isnon-affirmative, a second determination (816) is made to determinewhether the address corresponding to the remote request received in step(802) corresponds to an address in the pending write queue correspondingto a memory write request executed by instructions which are speculativerelative to other instructions, and speculative relative to the possibleoccurrence of a transaction abort or transaction rollback. If a responseto the determination at step (816) is affirmative, a hold is placed onthe read request until the request is determined to be non-speculative(818), e.g. until the address is removed from the queue. Following step(818), it is again determined if a non-speculative write indicator,e.g., the write set (1238) of FIG. 12, or the TX DIRTY indicator (1052)of FIG. 10, has been set in the cache by a local processor (820). Anaffirmative response at step (820) is proceeds to step (814) to initiaterecovery of the transaction due to a conflict. However, anon-affirmative response to the determinations at either step (816) or(820) is an indication that there is no interference associated with theread request. A cache line for the requests access is provided to therequesting processor (822), and the flow for checking for a conflictbetween a request from a remote processor and a currently activetransaction on a local processor concludes (824). As shown, for a readrequest, one or more over-indication tests are used to hold a responseto the received request until the speculative instruction is committedor flushed. In another embodiment, the over-indication test of step(820) can be performed by determining whether an instructioncorresponding to an interfering address pending queue entry is flushed(no recovery necessary at step (822)), or is committed and initiatingrecovery at step (814) responsive to a detected interference.

The determination for a potential conflict for a write request isdifferent from that of a read request. More specifically and as shown atstep (826), the write request tests for presence of interference withrespect to non-speculative instructions of an active transaction, andspeculative instructions of a transaction. At step (826), the methodperforms a determination for interference of the request of (802) witheither a non-speculative write set indicator, e.g. W set (1238) or TXDIRTY (1052), or a non-speculative read set indicator, e.g. R set (1232)or TX READ (1048), in the cache, corresponding to non-speculative memoryread or memory write instructions (including but not limited to load andstore instructions, respectively, and computational instructions havingat least a memory write and memory read, operand, respectively, thenon-speculative nature referring to the instruction beingnon-speculative with respect to other instructions, but beingspeculative with respect to the possibility of a transaction abort ortransaction rollback occurring). Following a non-affirmative response tothe determination at step (826) is following by a second determination(828) to determine whether the address corresponding to the remoterequest received in step (802) corresponds to an address in the pendingread queue or the pending write queue corresponding to a memory writerequest executed by instructions which are speculative relative to otherinstructions, and speculative relative to the possible occurrence of atransaction abort or transaction rollback. If a response to thedetermination at step (828) is affirmative, a hold is placed on thewrite request until the request is determined to be non-speculative(830), e.g. until the address is removed from the queue. Following step(830), the method performs a determination for interference (832) of therequest of (802) with either a non-speculative write set indicator, e.g.W set (1238) or TX DIRTY (1052), or a non-speculative read setindicator, e.g. R set (1232) or TX READ (1048), in the cache,corresponding to non-speculative memory read or memory writeinstructions. An affirmative response at step (832) proceeds to step(838) to initiate recovery of the transaction due to a conflict.However, a non-affirmative response to the determinations at either step(828) or (832) is an indication that there is no interference associatedwith the write request. A cache line for the requested access isprovided to the requesting processor (834), and the flow for checkingfor a conflict between a request from a remote processor and a currentlyactive transaction on a local processor concludes (836). As shown, for awrite request, one or more over-indication tests are used to hold aresponse to the received request until the speculative instruction iscommitted or flushed. In another embodiment, the over-indication test ofstep (832) can be performed by determining whether an instructioncorresponding to an interfering address pending queue entry is flushed(no recovery necessary at step (834)), or is committed and initiatingrecovery at step (838) responsive to a detected interference.

As shown and described in FIGS. 8A-8C, the transaction interferenceresolution determines a potential interference exclusively conflictingwith a speculative instruction. Based on the potential interference, oneor two over-indication tests are conducted, and in limited circumstancesa response to the requesting processor is placed on hold until thespeculative instruction is either committed or flushed. The act ofplacing the transaction on hold does not place the transaction in thisstate indefinitely, and effectively defers the response to therequesting processor, e.g. remote processor, and prevents the requestingprocessor from performing a transaction resolution.

Referring to FIGS. 9A-9C, a flow chart (900) is provided illustratinganother process for holding a response to a requesting processor basedon potential interference associated with a transaction. Similar to theprocess shown and described in FIGS. 1 and 8A-8C, a remote processorrequests data (902). In response to receiving such a request, logicoperatively coupled to a local processor accesses at least one localcache (904). It is then determined if the data is present in theaccessed local cache (906). If the data is not present, the flow ends(908). However, if the data is present, it is then ascertained if therequest received from the remote processor requests the data for a readrequest or a write request (910). Multiple transactions can read thesame data. To determine whether interference exits for a read request, aset of alternative tests is conducted. First, a test is conducted todetermine whether the request received in step (902) corresponds to aread or a write request. If the request received at step (902)corresponds to a read request, control transfers to step (912).Alternatively, if the request received at step (902) corresponds to awrite, control transfers to step (930).

In order to determine interference for a received read request, at step(912) it is determined if a non-speculative write indicator, e.g., thewrite set (1238) of FIG. 12, or the TX DIRTY indicator (1052) of FIG.10, has been set in the cache by a local processor. Presence of the bitis evidence that that requested data is the subject of a pending memorywrite request corresponding to a non-speculative instruction(non-speculative with respect to other instructions, but speculativewith respect to a transaction abort or rollback occurring) of atransaction on the present processor, indicating that a transaction onthe local processor may be in the process of changing the requested dataassociated with the transaction. An affirmative response to thedetermination at step (912) is followed by initiating recovery of thetransaction due to a conflict (914). If the response at step (912) isnon-affirmative, a second determination (916) is made to determinewhether the address corresponding to the remote request received in step(902) corresponds to an address in the pending write queue correspondingto a memory write request executed by instructions which are speculativerelative to other instructions, and speculative relative to the possibleoccurrence of a transaction abort or transaction rollback. If a responseto the determination at step (916) is affirmative, a deferred responseindication is sent to the requesting processor (918), and a hold isplaced on the read request until the request is determined to benon-speculative (920), e.g. until the address is removed from the queue.Following the hold at step (920), it is determined if a deferredresponse has been received (922). In one embodiment, the hold at step(920) is limited in time, and following the expiration of the limited,the test at step (922) is conducted. A deferred response at step (922)is an indication that there is a conflict, and recovery is initiated(914). However, if the deferred response has not been received at step(922), a second over indication test takes place to determine if anon-speculative write indicator, e.g., the write set (1238) of FIG. 12,or the TX DIRTY indicator (1052) of FIG. 10, has been set in the cacheby a local processor (924). An affirmative response at step (924) isproceeds to step (914) to initiate recovery of the transaction due to aconflict. However, a non-affirmative response to the determinations ateither step (916) or (924) is an indication that there is nointerference associated with the read request. A cache line for therequested access is provided to the requesting processor (926), and theflow for checking for a conflict between a request from a remoteprocessor and a currently active transaction on a local processorconcludes (928). Accordingly, for a read request, one or moreover-indication tests are used to hold a response to the receivedrequest until the speculative instruction is committed or flushed.

The determination for a potential conflict for a write request isdifferent from that of a read request. More specifically and as shown atstep (930), the write request tests for presence of interference withrespect to non-speculative instructions of an active transaction, andspeculative instructions of a transaction. At step (930), the methodperforms a determination for interference of the request of (902) witheither a non-speculative write set indicator, e.g. W set (1238) or TXDIRTY (1052), or a non-speculative read set indicator, e.g. R set (1232)or TX READ (1048), in the cache, corresponding to non-speculative memoryread or memory write instructions (including but not limited to load andstore instructions, respectively, and computational instructions havingat least a memory write and memory read, operand, respectively, thenon-speculative nature referring to the instruction beingnon-speculative with respect to other instructions, but beingspeculative with respect to the possibility of a transaction abort ortransaction rollback occurring). Following a non-affirmative response tothe determination at step (930) is following by a second determination(932) to determine whether the address corresponding to the remoterequest received in step (902) corresponds to an address in the pendingread queue or the pending write queue corresponding to a memory writerequest executed by instructions which are speculative relative to otherinstructions, and speculative relative to the possible occurrence of atransaction abort or transaction rollback. If a response to thedetermination at step (932) is affirmative, a deferred responseindication is sent to the requesting processor (934), and a hold isplaced on the write request until the request is determined to benon-speculative (936), e.g. until the address is removed from the queue.Following the hold at step (936), it is determined if a deferredresponse has been received (938). In one embodiment, the hold at step(936) is limited in time, and following the expiration of the limit, thetest at step (938) is conducted. A deferred response at step (938) is anindication that there is a conflict, and recovery is initiated (946).However, if the deferred response has not been received at step (938), asecond over indication test takes place to determine if anon-speculative write indicator, e.g., the write set (1238) of FIG. 12,or the TX DIRTY indicator (1052) of FIG. 10, has been set in the cacheby a local processor (940). An affirmative response at step (940)proceeds to step (946) to initiate recovery of the transaction due to aconflict. However, a non-affirmative response to the determinations ateither step (932) or (940) is an indication that there is nointerference associated with the read request. A cache line for therequested access is provided to the requesting processor (942), and theflow for checking for a conflict between a request from a remoteprocessor and a currently active transaction on a local processorconcludes (944). Accordingly, for a write request, one or moreover-indication tests are used to hold a response to the receivedrequest until the speculative instruction is committed or flushed.

The read set indicator pertains to reading new data versus old data. Thewrite set indicator tests and validates that two transactions cannotwrite to the same location. A non-affirmative response to theevaluation(s) at step (930) is followed by a first of two possibleover-indication tests associated with the write request evaluation. Thefirst test evaluates for presence of a write set indicator or a read setindicator in the cache (932), and a second over indication test takesplace to determine if the non-speculative write indicator ornon-speculative read indicator is present in the cache (940).

The process for holding a response to a requesting processor based onpotential interference associated with a transaction as shown anddescribed herein includes two steps in addition to the over-indicationtest. These two steps are present for both the read transactionprocessing and the write transaction processing. The steps includesending a deferred response indicator to the requesting processor, whicheffectively communicates to the processor that there will be a delay inprocessing.

The over-indication tests shown and described in FIGS. 8 and 9, delayinterference processing in response to an interference with aspeculative access, i.e., with an address in a pending access queue whenthe read and/or write indicator(s) in the cache are not present, so thatthe conflict is exclusive with the speculative access. These processesdo not just avoid interference, but take it further by determiningwhether the speculative instruction will complete. Then, if thespeculative instruction completes, the interference can be indicated andrecovery initiated, and if the speculative instruction does not completea non-interference can be indicated and the data provided to therequesting processor. In one embodiment, a deadlock detection isprovided, such as a time-out or a new protocol response, where anyprocessor that is deferring a response must not allow a load or storefor which it has received a deferring response to hold up instructioncommit since that creates a potential for deadlock. Instead, when aprocessor issues a deferring response in response to a remote request ithas received (steps (918), and (934)), it must not allow a request thatit has issued and for which it has received a deferred response fromanother processor to halt forward progress. Rather, optionally after aninitial time-out, i.e., after allowing for a lapse of time in which thesituation might resolve when a non-circular dependence is present, sucha processor must abort its own transaction. In a variety of embodiment,rules may be provided wherein between two dependent processes, only oneinitiate's recovery. In one embodiment, a processor with a higher threadID initiates recovery on a possible deadlock, and a processor with alower processor number does not. In another embodiment, lower/higher,even/odd, or token-based determinations are used (e.g., a process mayhave an “immunity token”.). In at least some embodiments, theseconditions are changed during operation of the system to balance systemload, establish fairness and increase system utilization.

With reference to FIG. 10, the IBM zEnterprise EC12 processor introducedthe transactional execution facility. The processor can decode 3instructions per clock cycle; simple instructions are dispatched assingle micro-ops, and more complex instructions are cracked intomultiple micro-ops. The micro-ops (Uops (1032 b)) are written into aunified issue queue (1016)), from where they can be issued out-of-order.Up to two fixed-point, one floating-point, two load/store, and twobranch instructions can execute every cycle. A Global Completion Table(GCT) (1032) holds every micro-op (1032 b) and a transaction nestingdepth (TND) (1032 a). The GCT (1032) is written in-order at decode time,tracks the execution status of each micro-op (1032 b), and completesinstructions when all micro-ops (1032 b) of the oldest instruction grouphave successfully executed.

The level 1 (L1) data cache (1040) is a 96 KB (kilo-byte) 6-wayassociative cache with 256 byte cache-lines and 4 cycle use latency,coupled to a private 1 MB (mega-byte) 8-way associative 2nd-level (L2)data cache (1068) with 7 cycles use-latency penalty for L1 (1040)misses. The L1 (1040) cache is the cache closest to a processor and Lncache is a cache at the nth level of caching. Both L1 (1040) and L2(1068) caches are store-through. Six cores on each central processor(CP) chip share a 48 MB 3rd-level store-in cache, and six CP chips areconnected to an off-chip 384 MB 4th-level cache, packaged together on aglass ceramic multi-chip module (MCM). Up to 4 multi-chip modules (MCMs)can be connected to a coherent symmetric multi-processor (SMP) systemwith up to 144 cores (not all cores are available to run customerworkload).

Coherency is managed with a variant of the MESI protocol. Cache-linescan be owned read-only (shared) or exclusive; the L1 (1040) and L2(1068) are store-through and thus do not contain dirty lines. The L3(1072) and L4 caches (not shown) are store-in and track dirty states.Each cache is inclusive of all its connected lower level caches.

Coherency requests are called “cross interrogates” (XI) and are senthierarchically from higher level to lower-level caches, and between theL4s. When one core misses the L1 (1040) and L2 (1068) and requests thecache line from its local L3 (1072), the L3 (1072) checks whether itowns the line, and if necessary sends an XI to the currently owning L2(1068)/L1 (1040) under that L3 (1072) to ensure coherency, before itreturns the cache line to the requestor. If the request also misses theL3 (1072), the L3 (1072) sends a request to the L4 (not shown), whichenforces coherency by sending XIs to all necessary L3s under that L4,and to the neighboring L4s. Then the L4 responds to the requesting L3which forwards the response to the L2 (1068)/L1 (1040).

Note that due to the inclusivity rule of the cache hierarchy, sometimescache lines are XI′ed from lower-level caches due to evictions onhigher-level caches caused by associativity overflows from requests toother cache lines. These XIs can be called “LRU XIs”, where LRU standsfor least recently used.

Making reference to yet another type of XI requests, Demote-XIstransition cache-ownership from exclusive into read-only state, andExclusive-XIs transition cache ownership from exclusive into invalidstate. Demote-XIs and Exclusive-XIs need a response back to the XIsender. The target cache can “accept” the XI, or send a “reject”response if it first needs to evict dirty data before accepting the XI.The L1 (1040)/L2 (1068) caches are store through, but may rejectdemote-XIs and exclusive XIs if they have stores in their store queuesthat need to be sent to L3 before downgrading the exclusive state. Arejected XI will be repeated by the sender. Read-only-XIs are sent tocaches that own the line read-only; no response is needed for such XIssince they cannot be rejected. The details of the SMP protocol aresimilar to those described for the IBM z10 by P. Mak, C. Walters, and G.Strait, in “IBM System z10 processor cache subsystem microarchitecture”,IBM Journal of Research and Development, Vol 53:1, 2009, which isincorporated by reference herein in its entirety.

Transactional Instruction Execution

FIG. 10 depicts example components of an example transactional executionenvironment, including a CPU and caches/components with which itinteracts. The instruction decode unit (1008) (IDU) keeps track of thecurrent transaction nesting depth (1012) (TND). When the IDU (1008)receives a TBEGIN instruction (1004), the nesting depth (1012) isincremented, and conversely decremented on TEND instructions. Thenesting depth (1012) is written into the GCT (1032) for every dispatchedinstruction. When a TBEGIN or TEND is decoded on a speculative path thatlater gets flushed, the IDU's (1008) nesting depth (1012) is refreshedfrom the youngest GCT (1032) entry that is not flushed. Thetransactional state is also written into the issue queue (1016) forconsumption by the execution units, mostly by the Load/Store Unit (LSU)(1080), which also has an effective address calculator (1036) includedin the LSU (1080). The TBEGIN instruction may specify a transactiondiagnostic block (TDB) for recording status information, should thetransaction abort before reaching a TEND instruction. In addition, theLoad/Store Unit (1080) is shown with two queues (1082) and (1084),referring the pending read queue and pending write queue respectively.The functionality of these queues is described above in FIGS. 1-9. Inone embodiment, the queues (1082) and (1084) may be combined into asingle queue while maintaining the functionality to support transactionexecution.

Similar to the nesting depth, the IDU (1008)/GCT (1032) collaborativelytrack the access register/floating-point register (AR/FPR) modificationmasks through the transaction nest; the IDU (1008) can place an abortrequest into the GCT (1032) when an AR/FPR-modifying instruction isdecoded and the modification mask blocks that. When the instructionbecomes next-to-complete, completion is blocked and the transactionaborts. Other restricted instructions are handled similarly, includingTBEGIN if decoded while in a constrained transaction, or exceeding themaximum nesting depth.

An outermost TBEGIN is cracked into multiple micro-ops depending on theGR-Save-Mask; each micro-op (1032 b) (including, for example μuop 0, μop1, and μop2) will be executed by one of the two fixed point units (FXUs)(1020) to save a pair of GRs (1028) into a special transaction-backupregister file (1024), that is used to later restore the GR (1028)content in case of a transaction abort. Also the TBEGIN spawns micro-ops(1032 b) to perform an accessibility test for the TDB if one isspecified; the address is saved in a special purpose register for laterusage in the abort case. At the decoding of an outermost TBEGIN, theinstruction address and the instruction text of the TBEGIN are alsosaved in special purpose registers for a potential abort processinglater on.

TEND and NTSTG are single micro-op (1032 b) instructions; NTSTG(non-transactional store) is handled like a normal store except that itis marked as non-transactional in the issue queue (1016) so that the LSU(1080) can treat it appropriately. TEND is a no-op at execution time,the ending of the transaction is performed when TEND completes.

As mentioned, instructions that are within a transaction are marked assuch in the issue queue (1016), but otherwise execute mostly unchanged;the LSU (1080) performs isolation tracking as described in the nextsection.

Since decoding is in-order, and since the IDU (1008) keeps track of thecurrent transactional state and writes it into the issue queue (1016)along with every instruction from the transaction, execution of TBEGIN,TEND, and instructions before, within, and after the transaction can beperformed out-of order. It is even possible (though unlikely) that TENDis executed first, then the entire transaction, and lastly the TBEGINexecutes. Program order is restored through the GCT (1032) at completiontime. The length of transactions is not limited by the size of the GCT(1032), since general purpose registers (GRs) (1028) can be restoredfrom the backup register file (1024).

During execution, the program event recording (PER) events are filteredbased on the Event Suppression Control, and a PER TEND event is detectedif enabled. Similarly, while in transactional mode, a pseudo-randomgenerator may be causing the random aborts as enabled by the TransactionDiagnostics Control.

Tracking for Transactional Isolation

The Load/Store Unit (1080) tracks cache lines that were accessed duringtransactional execution, and triggers an abort if an XI from another CPU(or an LRU-XI) conflicts with the footprint. If the conflicting XI is anexclusive or demote XI, the LSU (1080) rejects the XI back to the L3(1072) in the hope of finishing the transaction before the L3 (1072)repeats the XI. This “stiff-arming” is very efficient in highlycontended transactions. In order to prevent hangs when two CPUsstiff-arm each other, a XI-reject counter is implemented, which triggersa transaction abort when a threshold is met.

The L1 cache directory (1040) is traditionally implemented with staticrandom access memories (SRAMs). For the transactional memoryimplementation, the valid bits 244 (64 rows×6 ways) of the directoryhave been moved into normal logic latches, and are supplemented with twomore bits per cache line: the TX-read 248 and TX-dirty 252 bits.

The TX-read (1048) bits are reset when a new outermost TBEGIN is decoded(which is interlocked against a prior still pending transaction). TheTX-read (1048) bit is set at execution time by every load instructionthat is marked “transactional” in the issue queue. Note that this canlead to over-marking if speculative loads are executed, for example on amispredicted branch path. The alternative of setting the TX-read (1048)bit at load completion time was too expensive for silicon area, sincemultiple loads can complete at the same time, requiring many read-portson the load-queue.

Stores execute the same way as in non-transactional mode, but atransaction mark is placed in the store queue (STQ) (1060) entry of thestore instruction. At write-back time, when the data from the STQ (1060)is written into the L1 (1040), the TX-dirty bit (1052) in theL1-directory (1056) is set for the written cache line. Store write-backinto the L1 (1040) occurs only after the store instruction hascompleted, and at most one store is written back per cycle. Beforecompletion and write-back, loads can access the data from the STQ (1060)by means of store-forwarding; after write-back, the CPU (Not Shown) canaccess the speculatively updated data in the L1 (1040). If thetransaction ends successfully, the TX-dirty bits (1052) of allcache-lines are cleared, and also the TX-marks of not yet written storesare cleared in the STQ (1060), effectively turning the pending storesinto normal stores.

On a transaction abort, all pending transactional stores are invalidatedfrom the STQ (1060), even those already completed. All cache lines thatwere modified by the transaction in the L1 (1040), that is, have theTX-dirty bit (1052) on, have their valid bits turned off, effectivelyremoving them from the L1 (1040) cache instantaneously.

The architecture requires that before completing a new instruction, theisolation of the transaction read- and write-set is maintained. Thisisolation is ensured by stalling instruction completion at appropriatetimes when XIs are pending; speculative out-of order execution isallowed, optimistically assuming that the pending XIs are to differentaddresses and not actually cause a transaction conflict. This designfits very naturally with the XI-vs-completion interlocks that areimplemented on prior systems to ensure the strong memory ordering thatthe architecture requires.

When the L1 (1040) receives an XI, L1 (1040) accesses the directory tocheck validity of the XI′ed address in the L1 (1040), and if the TX-readbit (1048) is active on the XI′ed line and the XI is not rejected, theLSU (1080) triggers an abort. When a cache line with active TX-read bit(1048) is LRU′ed from the L1 (1040), a special LRU-extension vectorremembers for each of the 64 rows of the L1 (1040) that a TX-read lineexisted on that row. Since no precise address tracking exists for theLRU extensions, any non-rejected XI that hits a valid extension row theLSU (1080) triggers an abort. Providing the LRU-extension effectivelyincreases the read footprint capability from the L1-size to the L2-sizeand associativity, provided no conflicts with other CPUs against thenon-precise LRU-extension tracking causes aborts.

The store footprint is limited by the store cache size and thusimplicitly by the L2 (1068) size and associativity. No LRU-extensionaction needs to be performed when a TX-dirty (1052) cache line is LRU′edfrom the L1 (1040).

In prior systems, since the L1 (1040) and L2 (1068) are store-throughcaches, every store instruction causes an L3 (1072) store access; withnow 6 cores per L3 (1072) and further improved performance of each core,the store rate for the L3 (1072) (and to a lesser extent for the L2(1068) becomes problematic for certain workloads. In order to avoidstore queuing delays, a gathering store cache (1064) had to be added,that combines stores to neighboring addresses before sending them to theL3 (1072).

For transactional memory performance, it is acceptable to invalidateevery TX-dirty (1052) cache line from the L1 (1040) on transactionaborts, because the L2 (1068) cache is very close (7 cycles L1 1040 misspenalty) to bring back the clean lines. However, it would beunacceptable for performance (and silicon area for tracking) to havetransactional stores write the L2 (1068) before the transaction ends andthen invalidate all dirty L2 (1068) cache lines on abort (or even worseon the shared L3 (1072)).

The two problems of store bandwidth and transactional memory storehandling can both be addressed with the gathering store cache (1064).The cache (1064) is a circular queue of 64 entries, each entry holding128 bytes of data with byte-precise valid bits. In non-transactionaloperation, when a store is received from the LSU (1080), the store cache(1064) checks whether an entry exists for the same address, and if sogathers the new store into the existing entry. If no entry exists, a newentry is written into the queue, and if the number of free entries fallsunder a threshold, the oldest entries are written back to the L2 (1068)and L3 (1072) caches.

When a new outermost transaction begins, all existing entries in thestore cache are marked closed so that no new stores can be gathered intothem, and eviction of those entries to L2 (1068) and L3 (1072) isstarted. From that point on, the transactional stores coming out of theLSU (1080) STQ (1060) allocate new entries, or gather into existingtransactional entries. The write-back of those stores into L2 (1068) andL3 (1072) is blocked, until the transaction ends successfully; at thatpoint subsequent (post-transaction) stores can continue to gather intoexisting entries, until the next transaction closes those entries again.

The store cache (1064) is queried on every exclusive or demote XI, andcauses an XI reject if the XI compares to any active entry. If the coreis not completing further instructions while continuously rejecting XIs,the transaction is aborted at a certain threshold to avoid hangs.

The LSU (1080) requests a transaction abort when the store cache (1064)overflows. The LSU (1080) detects this condition when it tries to send anew store that cannot merge into an existing entry, and the entire storecache (1064) is filled with stores from the current transaction. Thestore cache (1064) is managed as a subset of the L2 (1068): whiletransactionally dirty lines can be evicted from the L1 (1040), they haveto stay resident in the L2 (1068) throughout the transaction. Themaximum store footprint is thus limited to the store cache size of64×128 bytes, and it is also limited by the associativity of the L2(1068). Since the L2 (1068) is 8-way associative and has 512 rows, it istypically large enough to not cause transaction aborts.

If a transaction aborts, the store cache (1064) is notified and allentries holding transactional data are invalidated. The store cache(1064) also has a mark per doubleword (8 bytes) whether the entry waswritten by a NTSTG instruction—those doublewords stay valid acrosstransaction aborts.

Traditionally, IBM mainframe server processors contain a layer offirmware called millicode which performs complex functions like certainCISC instruction executions, interruption handling, systemsynchronization, and RAS. Millicode includes machine dependentinstructions as well as instructions of the instruction set architecture(ISA) that are fetched and executed from memory similarly toinstructions of application programs and the operating system (OS).Firmware resides in a restricted area of main memory that customerprograms cannot access. When hardware detects a situation that needs toinvoke millicode, the instruction fetching unit 204 switches into“millicode mode” and starts fetching at the appropriate location in themillicode memory area Millicode may be fetched and executed in the sameway as instructions of the instruction set architecture (ISA), and mayinclude ISA instructions.

For transactional memory, millicode is involved in various complexsituations. Every transaction abort invokes a dedicated millicodesub-routine to perform the necessary abort steps. The transaction-abortmillicode starts by reading special-purpose registers (SPRs) holding thehardware internal abort reason, potential exception reasons, and theaborted instruction address, which millicode then uses to store a TDB ifone is specified. The TBEGIN instruction text is loaded from an SPR toobtain the GR-save-mask, which is needed for millicode to know which GRs(1038) to restore.

The CPU supports a special millicode-only instruction to read out thebackup-GRs (1024) and copy them into the main GRs (1028). The TBEGINinstruction address is also loaded from an SPR to set the newinstruction address in the PSW to continue execution after the TBEGINonce the millicode abort sub-routine finishes. That PSW may later besaved as program-old PSW in case the abort is caused by a non-filteredprogram interruption.

The TABORT instruction may be millicode implemented; when the IDU (1208)decodes TABORT, it instructs the instruction fetch unit to branch intoTABORT's millicode, from which millicode branches into the common abortsub-routine.

The Extract Transaction Nesting Depth (ETND) instruction may also bemillicoded, since it is not performance critical; millicode loads thecurrent nesting depth out of a special hardware register and places itinto a GR (1028). The PPA instruction is millicoded; it performs theoptimal delay based on the current abort count provided by software asan operand to PPA, and also based on other hardware internal state.

For constrained transactions, millicode may keep track of the number ofaborts. The counter is reset to 0 on successful TEND completion, or ifan interruption into the OS occurs (since it is not known if or when theOS will return to the program). Depending on the current abort count,millicode can invoke certain mechanisms to improve the chance of successfor the subsequent transaction retry. The mechanisms involve, forexample, successively increasing random delays between retries, andreducing the amount of speculative execution to avoid encounteringaborts caused by speculative accesses to data that the transaction isnot actually using. As a last resort, millicode can broadcast to otherCPUs to stop all conflicting work, retry the local transaction, beforereleasing the other CPUs to continue normal processing. Multiple CPUsmust be coordinated to not cause deadlocks, so some serializationbetween millicode instances on different CPUs is required.

In today's systems, accurate tracking of transactional read and writesets is difficult if not impossible. When a possibly speculative readaccess is made, a cache line is indicated to be in the read set. When anevent causing discarding of speculative execution occurs, no resetoccurs. This is similar, for write sets. Thus, read and write sets fortransactions necessarily contain speculative over indication, when alater branch mis-prediction recovery has in fact ejected that (read orwrite) access from the (read or write) set. Therefore, it may beadvantageous to improve the tracking of read and write sets with thepending read and write queues (1082) and (1084), respectively,associated with transactions themselves and allow a processor to recoverread and write sets that have been unnecessarily augmented withspeculative addresses when a mis-speculation is discovered. As such,according to at least one embodiment.

The components of an example CPU described above in FIG. 10 has beenlabeled with tools in the form of read and write queues (1080) and(1082). The tools are implemented to manage speculative load and storeinstructions for tracking processor transaction read and write sets andeliminating speculative mis-predictions. The functionality of the readand write queues (1080) and (1082) for managing transactional executionare shown and described in the flow charts shown in FIGS. 1-9. Thequeues (1080) and (1082) may be implemented in programmable hardwaredevices such as field programmable gate arrays, programmable arraylogic, programmable logic devices, or the like. The queues (1080) and(1082) may also be implemented in software for execution by varioustypes of processors. An identified functional unit of executable codemay, for instance, comprise one or more physical or logical blocks ofcomputer instructions which may, for instance, be organized as anobject, procedure, function, or other construct. Nevertheless, theexecutable of the queues (1080) and (1082) need not be physicallylocated together, but may comprise disparate instructions stored indifferent locations which, when joined logically together, comprise thetools and achieve the stated purpose of the tool.

Indeed, executable code could be a single instruction, or manyinstructions, and may even be distributed over several different codesegments, among different applications, and across several memorydevices. Similarly, operational data may be identified and illustratedherein within the tool, and may be embodied in any suitable form andorganized within any suitable type of data structure. The operationaldata may be collected as a single data set, or may be distributed overdifferent locations including over different storage devices, and mayexist, at least partially, as electronic signals on a system or network.

Furthermore, the described features, structures, or characteristics maybe combined in any suitable manner in one or more embodiments. In thefollowing description, numerous specific details are provided, such asexamples of agents, to provide a thorough understanding of theembodiment(s). One skilled in the relevant art will recognize, however,that the embodiment(s) can be practiced without one or more of thespecific details, or with other methods, components, materials, etc. Inother instances, well-known structures, materials, or operations are notshown or described in detail to avoid obscuring aspects.

The present embodiment(s) may be a system, a method, and/or a computerprogram product. The computer program product may include a computerreadable storage medium (or media) having computer readable programinstructions thereon for causing a processor to carry out aspects of thepresent embodiment(s).

The computer readable storage medium can be a tangible device that canretain and store instructions for use by an instruction executiondevice. The computer readable storage medium may be, for example, but isnot limited to, an electronic storage device, a magnetic storage device,an optical storage device, an electromagnetic storage device, asemiconductor storage device, or any suitable combination of theforegoing. A non-exhaustive list of more specific examples of thecomputer readable storage medium includes the following: a portablecomputer diskette, a hard disk, a random access memory (RAM), aread-only memory (ROM), an erasable programmable read-only memory (EPROMor Flash memory), a static random access memory (SRAM), a portablecompact disc read-only memory (CD-ROM), a digital versatile disk (DVD),a memory stick, a floppy disk, a mechanically encoded device such aspunch-cards or raised structures in a groove having instructionsrecorded thereon, and any suitable combination of the foregoing. Acomputer readable storage medium, as used herein, is not to be construedas being transitory signals per se, such as radio waves or other freelypropagating electromagnetic waves, electromagnetic waves propagatingthrough a waveguide or other transmission media (e.g., light pulsespassing through a fiber-optic cable), or electrical signals transmittedthrough a wire.

Computer readable program instructions described herein can bedownloaded to respective computing/processing devices from a computerreadable storage medium or to an external computer or external storagedevice via a network, for example, the Internet, a local area network, awide area network and/or a wireless network. The network may comprisecopper transmission cables, optical transmission fibers, wirelesstransmission, routers, firewalls, switches, gateway computers and/oredge servers. A network adapter card or network interface in eachcomputing/processing device receives computer readable programinstructions from the network and forwards the computer readable programinstructions for storage in a computer readable storage medium withinthe respective computing/processing device.

Computer readable program instructions for carrying out operations ofthe present embodiment(s) may be assembler instructions,instruction-set-architecture (ISA) instructions, machine instructions,machine dependent instructions, microcode, firmware instructions,state-setting data, or either source code or object code written in anycombination of one or more programming languages, including an objectoriented programming language such as Smalltalk, C++ or the like, andconventional procedural programming languages, such as the “C”programming language or similar programming languages. The computerreadable program instructions may execute entirely on the user'scomputer, partly on the user's computer, as a stand-alone softwarepackage, partly on the user's computer and partly on a remote computeror entirely on the remote computer or server. In the latter scenario,the remote computer may be connected to the user's computer through anytype of network, including a local area network (LAN) or a wide areanetwork (WAN), or the connection may be made to an external computer(for example, through the Internet using an Internet Service Provider).In some embodiments, electronic circuitry including, for example,programmable logic circuitry, field-programmable gate arrays (FPGA), orprogrammable logic arrays (PLA) may execute the computer readableprogram instructions by utilizing state information of the computerreadable program instructions to personalize the electronic circuitry,in order to perform aspects of the present embodiment(s).

Aspects of the present embodiment(s) are described herein with referenceto flowchart illustrations and/or block diagrams of methods, apparatus(systems), and computer program products. It will be understood thateach block of the flowchart illustrations and/or block diagrams, andcombinations of blocks in the flowchart illustrations and/or blockdiagrams, can be implemented by computer readable program instructions.

These computer readable program instructions may be provided to aprocessor of a general purpose computer, special purpose computer, orother programmable data processing apparatus to produce a machine, suchthat the instructions, which execute via the processor of the computeror other programmable data processing apparatus, create means forimplementing the functions/acts specified in the flowchart and/or blockdiagram block or blocks. These computer readable program instructionsmay also be stored in a computer readable storage medium that can directa computer, a programmable data processing apparatus, and/or otherdevices to function in a particular manner, such that the computerreadable storage medium having instructions stored therein comprises anarticle of manufacture including instructions which implement aspects ofthe function/act specified in the flowchart and/or block diagram blockor blocks.

The computer readable program instructions may also be loaded onto acomputer, other programmable data processing apparatus, or other deviceto cause a series of operational steps to be performed on the computer,other programmable apparatus or other device to produce a computerimplemented process, such that the instructions which execute on thecomputer, other programmable apparatus, or other device implement thefunctions/acts specified in the flowchart and/or block diagram block orblocks.

The flowchart and block diagrams in the Figures illustrate thearchitecture, functionality, and operation of possible implementationsof systems, methods, and computer program products according to thevarious embodiments. In this regard, each block in the flowchart orblock diagrams may represent a module, segment, or portion ofinstructions, which comprises one or more executable instructions forimplementing the specified logical function(s). In some alternativeimplementations, the functions noted in the block may occur out of theorder noted in the figures. For example, two blocks shown in successionmay, in fact, be executed substantially concurrently, or the blocks maysometimes be executed in the reverse order, depending upon thefunctionality involved. It will also be noted that each block of theblock diagrams and/or flowchart illustration, and combinations of blocksin the block diagrams and/or flowchart illustration, can be implementedby special purpose hardware-based systems that perform the specifiedfunctions or acts or carry out combinations of special purpose hardwareand computer instructions.

The terminology used herein is for the purpose of describing particularembodiments only and is not intended to be limiting. As used herein, thesingular forms “a”, “an” and “the” are intended to include the pluralforms as well, unless the context clearly indicates otherwise. It willbe further understood that the terms “comprises” and/or “comprising,”when used in this specification, specify the presence of statedfeatures, integers, steps, operations, elements, and/or components, butdo not preclude the presence or addition of one or more other features,integers, steps, operations, elements, components, and/or groupsthereof.

The corresponding structures, materials, acts, and equivalents of allmeans or step plus function elements in the claims below are intended toinclude any structure, material, or act for performing the function incombination with other claimed elements as specifically claimed. Thedescription of the present embodiment(s) has been presented for purposesof illustration and description, but is not intended to be exhaustive orlimited to the embodiment(s) in the form disclosed. Many modificationsand variations will be apparent to those of ordinary skill in the artwithout departing from the scope and spirit of the embodiment(s). Theembodiment was chosen and described in order to best explain theprinciples and the practical application, and to enable others ofordinary skill in the art to understand the embodiment(s) for variousembodiments with various modifications as are suited to the particularuse contemplated. Accordingly, the implementation of the read and writequeues corresponding to speculatively executed read and writeinstructions, respectively, function to resolve transactioninterference.

It will be appreciated that, although specific embodiments have beendescribed herein for purposes of illustration, various modifications maybe made without departing from the spirit and scope of the invention. Inparticular, different coherence protocols, coherence granules, trackinggranularity of transactional read and write sets other than a cache linemay be used. Similarly, but or directory based protocols may be used tocommunicate. Accordingly, the scope of protection of this invention islimited only by the following claims and their equivalents.

We claim:
 1. A computer program product for tracking processortransactional read and write sets, the computer program productcomprising a computer readable storage medium having program codeembodied therewith, the program code executable by a processing unit to:maintain a non-speculative read set indication of non-speculative readsof data stored in a cache by a processor cache unit for a transaction bya first requestor; maintain a non-speculative write set indication ofnon-speculative writes of written data stored in the cache by theprocessor cache unit by the first requestor; maintain a first queue ofaddresses corresponding to speculatively executed load memory readinstructions corresponding to speculative members of a read set;maintain a second queue of addresses corresponding to speculativelyexecuted store memory write instructions corresponding to speculativemembers of a write set; perform a transaction interference resolutionresponsive to receiving a request for data by a remote processor,including determining a potential transaction interference exclusivelyconflicting with a speculative instruction; and hold a response to thereceived request until the speculative instruction is in a stateselected from the group consisting of: committed and flushed.
 2. Thecomputer program product of claim 1, further comprising program codeexecutable by the processing unit to determine if the speculativeinstruction will complete.
 3. The computer program product of claim 2,further comprising program code to determine if the speculativeinstruction completes, and to indicate interference and initiaterecovery.
 4. The computer program product of claim 2, further comprisingto determine if the speculative instruction does not complete, and toindicate non-interference and provide a cache line for the requestedaccess to the remote processor.
 5. The computer program product of claim1, further comprising program code to notify the remote processor thatthe response is being deferred, and advance the response in the remoteprocessor responsive to an interference with speculative instructions ofthe remote processor when the remote processor has received a deferredresponse from a local processor.
 6. The computer program product ofclaim 1, wherein the received request is a write request, and furthercomprising program code in support of the transaction interference toadditionally check the first queue and the non-speculative read set forthe pending address.
 7. The computer program product of claim 1, furthercomprising program code to determine that a deadlock may be possible,and responsive to the determination, indicate interference and initiaterecovery.
 8. The computer program product of claim 7, wherein thedetermination is based on one of exceeding a fixed number of cyclessince one of a decision to hold and a previous instruction completion.9. A computer system comprising: a processing unit operatively coupledto memory; a tool in communication with the processing unit to trackprocessor transactional read and write sets, including: maintain anon-speculative read set indication of non-speculative reads of datastored in a cache by a processor cache unit for a transaction by a firstrequestor; maintain a non-speculative write set indication ofnon-speculative writes of written data stored in the cache by theprocessor cache unit by the first requestor; maintain a first queue ofaddresses corresponding to speculatively executed load memory readinstructions corresponding to speculative members of a read set;maintain a second queue of addresses corresponding to speculativelyexecuted store memory write instructions corresponding to speculativemembers of a write set; perform a transaction interference resolutionresponsive to receiving a request for data by a remote processor,including determining a potential transaction interference exclusivelyconflicting with a speculative instruction; and hold a response to thereceived request until the speculative instruction is in a stateselected from the group consisting of: committed and flushed.
 10. Thesystem of claim 9, further comprising the tool to determine if thespeculative instruction will complete.
 11. The system of claim 10,further comprising the tool to determine completion of the speculativeinstruction, and to respond to the received request including indicateinterference with the speculative instruction and initiate recovery. 12.The system of claim 10, further comprising the tool to determinenon-completion of the speculative instruction, and respond to thereceived request including to indicate non-interference and provide acache line for the requested access to the remote processor.
 13. Thesystem of claim 9, further comprising the tool to notify the remoteprocessor that the response is being deferred, and to advance theresponse in the remote processor responsive to an interference withspeculative instructions of the remote processor when the remoteprocessor has received a deferred response from a local processor. 14.The system of claim 9, wherein the received request is a write request,and further comprising program code in support of the transactioninterference to additionally check the first queue and thenon-speculative read set for the pending address.
 15. The system ofclaim 9, further comprising the tool to determine a deadlock, andresponsive to the determination, indicate interference and initiaterecovery.
 16. The system of claim 15, wherein the determination of thedeadlock is based on one of exceeding a fixed number of cycles since adecision to hold and exceeding a fixed number of cycles since a previousinstruction completion.